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By submitting this Internet-Draft, each author represents that any applicable patent or other IPR claims of which he or she is aware have been or will be disclosed, and any of which he or she becomes aware will be disclosed, in accordance with Section 6 of BCP 79.
Internet-Drafts are working documents of the Internet Engineering Task Force (IETF), its areas, and its working groups. Note that other groups may also distribute working documents as Internet-Drafts.
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The list of Internet-Draft Shadow Directories can be accessed at http://www.ietf.org/shadow.html.
This Internet-Draft will expire on April 25, 2007.
Copyright © The Internet Society (2006).
This Internet-Draft describes NFSv4 minor version one, including features retained from the base protocol and protocol extensions made subsequently. The current draft includes description of the major extensions, Sessions, Directory Delegations, and parallel NFS (pNFS). This Internet-Draft is an active work item of the NFSv4 working group. Active and resolved issues may be found in the issue tracker at: http://www.nfsv4-editor.org/cgi-bin/roundup/nfsv4. New issues related to this document should be raised with the NFSv4 Working Group nfsv4@ietf.org and logged in the issue tracker.
The key words "MUST", "MUST NOT", "REQUIRED", "SHALL", "SHALL NOT", "SHOULD", "SHOULD NOT", "RECOMMENDED", "MAY", and "OPTIONAL" in this document are to be interpreted as described in RFC 2119 (Bradner, S., “Key words for use in RFCs to Indicate Requirement Levels,” March 1997.) [1].
1.
Introduction
1.1.
The NFSv4.1 Protocol
1.2.
NFS Version 4 Goals
1.3.
Minor Version 1 Goals
1.4.
Overview of NFS version 4.1 Features
1.4.1.
RPC and Security
1.4.2.
Protocol Structure
1.4.3.
File System Model
1.4.4.
Locking Facilities
1.5.
General Definitions
1.6.
Differences from NFSv4.0
2.
Core Infrastructure
2.1.
Introduction
2.2.
RPC and XDR
2.2.1.
RPC-based Security
2.3.
COMPOUND and CB_COMPOUND
2.4.
Client Identifiers
2.4.1.
Server Release of Clientid
2.5.
Security Service Negotiation
2.5.1.
NFSv4 Security Tuples
2.5.2.
SECINFO and SECINFO_NO_NAME
2.5.3.
Security Error
2.6.
Minor Versioning
2.7.
Non-RPC-based Security Services
2.7.1.
Authorization
2.7.2.
Auditing
2.7.3.
Intrusion Detection
2.8.
Transport Layers
2.8.1.
Required and Recommended Properties of Transports
2.8.2.
Client and Server Transport Behavior
2.8.3.
Ports
2.9.
Session
2.9.1.
Motivation and Overview
2.9.2.
NFSv4 Integration
2.9.3.
Channels
2.9.4.
Exactly Once Semantics
2.9.5.
RDMA Considerations
2.9.6.
Sessions Security
2.9.7.
Session Mechanics - Steady State
2.9.8.
Session Mechanics - Recovery
3.
Protocol Data Types
3.1.
Basic Data Types
3.2.
Structured Data Types
4.
Filehandles
4.1.
Obtaining the First Filehandle
4.1.1.
Root Filehandle
4.1.2.
Public Filehandle
4.2.
Filehandle Types
4.2.1.
General Properties of a Filehandle
4.2.2.
Persistent Filehandle
4.2.3.
Volatile Filehandle
4.3.
One Method of Constructing a Volatile Filehandle
4.4.
Client Recovery from Filehandle Expiration
5.
File Attributes
5.1.
Mandatory Attributes
5.2.
Recommended Attributes
5.3.
Named Attributes
5.4.
Classification of Attributes
5.5.
Mandatory Attributes - Definitions
5.6.
Recommended Attributes - Definitions
5.7.
Time Access
5.8.
Interpreting owner and owner_group
5.9.
Character Case Attributes
5.10.
Quota Attributes
5.11.
mounted_on_fileid
5.12.
send_impl_id and recv_impl_id
5.13.
fs_layout_type
5.14.
layout_type
5.15.
layout_hint
5.16.
mdsthreshold
5.17.
Retention Attributes
6.
Access Control Lists
6.1.
Goals
6.2.
File Attributes Discussion
6.2.1.
ACL Attribute
6.2.2.
mode Attribute
6.3.
Common Methods
6.3.1.
Interpreting an ACL
6.3.2.
Computing a Mode Attribute from an ACL
6.4.
Requirements
6.4.1.
Setting the mode and/or ACL Attributes
6.4.2.
Retrieving the mode and/or ACL Attributes
6.4.3.
Creating New Objects
7.
Single-server Name Space
7.1.
Server Exports
7.2.
Browsing Exports
7.3.
Server Pseudo File System
7.4.
Multiple Roots
7.5.
Filehandle Volatility
7.6.
Exported Root
7.7.
Mount Point Crossing
7.8.
Security Policy and Name Space Presentation
8.
File Locking and Share Reservations
8.1.
Locking
8.1.1.
Client and Session ID
8.1.2.
State-owner and Stateid Definition
8.1.3.
Use of the Stateid and Locking
8.2.
Lock Ranges
8.3.
Upgrading and Downgrading Locks
8.4.
Blocking Locks
8.5.
Lease Renewal
8.6.
Crash Recovery
8.6.1.
Client Failure and Recovery
8.6.2.
Server Failure and Recovery
8.6.3.
Network Partitions and Recovery
8.7.
Server Revocation of Locks
8.8.
Share Reservations
8.9.
OPEN/CLOSE Operations
8.10.
Open Upgrade and Downgrade
8.11.
Short and Long Leases
8.12.
Clocks, Propagation Delay, and Calculating Lease Expiration
8.13.
Vestigial Locking Infrastructure From V4.0
9.
Client-Side Caching
9.1.
Performance Challenges for Client-Side Caching
9.2.
Delegation and Callbacks
9.2.1.
Delegation Recovery
9.3.
Data Caching
9.3.1.
Data Caching and OPENs
9.3.2.
Data Caching and File Locking
9.3.3.
Data Caching and Mandatory File Locking
9.3.4.
Data Caching and File Identity
9.4.
Open Delegation
9.4.1.
Open Delegation and Data Caching
9.4.2.
Open Delegation and File Locks
9.4.3.
Handling of CB_GETATTR
9.4.4.
Recall of Open Delegation
9.4.5.
Clients that Fail to Honor Delegation Recalls
9.4.6.
Delegation Revocation
9.5.
Data Caching and Revocation
9.5.1.
Revocation Recovery for Write Open Delegation
9.6.
Attribute Caching
9.7.
Data and Metadata Caching and Memory Mapped Files
9.8.
Name Caching
9.9.
Directory Caching
10.
Multi-server Name Space
10.1.
Location attributes
10.2.
File System Presence or Absence
10.3.
Getting Attributes for an Absent File System
10.3.1.
GETATTR Within an Absent File System
10.3.2.
READDIR and Absent File Systems
10.4.
Uses of Location Information
10.4.1.
File System Replication
10.4.2.
File System Migration
10.4.3.
Referrals
10.5.
Additional Client-side Considerations
10.6.
Effecting File System Transitions
10.6.1.
File System Transitions and Simultaneous Access
10.6.2.
Simultaneous Use and Transparent Transitions
10.6.3.
Filehandles and File System Transitions
10.6.4.
Fileid's and File System Transitions
10.6.5.
Fsid's and File System Transitions
10.6.6.
The Change Attribute and File System Transitions
10.6.7.
Lock State and File System Transitions
10.6.8.
Write Verifiers and File System Transitions
10.7.
Effecting File System Referrals
10.7.1.
Referral Example (LOOKUP)
10.7.2.
Referral Example (READDIR)
10.8.
The Attribute fs_absent
10.9.
The Attribute fs_locations
10.10.
The Attribute fs_locations_info
10.10.1.
The location4_server Structure
10.10.2.
The location4_info Structure
10.10.3.
The location4_item Structure
10.11.
The Attribute fs_status
11.
Directory Delegations
11.1.
Introduction to Directory Delegations
11.2.
Directory Delegation Design (in brief)
11.3.
Recommended Attributes in support of Directory Delegations
11.4.
Delegation Recall
11.5.
Directory Delegation Recovery
12.
Parallel NFS (pNFS)
12.1.
Introduction
12.2.
General Definitions
12.2.1.
Metadata Server
12.2.2.
Client
12.2.3.
Storage Device
12.2.4.
Storage Protocol
12.2.5.
Control Protocol
12.2.6.
Metadata
12.2.7.
Layout
12.3.
pNFS protocol semantics
12.3.1.
Definitions
12.3.2.
Guarantees Provided by Layouts
12.3.3.
Getting a Layout
12.3.4.
Committing a Layout
12.3.5.
Recalling a Layout
12.3.6.
Metadata Server Write Propagation
12.3.7.
Crash Recovery
12.3.8.
Security Considerations
12.4.
The NFSv4.1 File Layout Type
12.4.1.
Session Considerations
12.4.2.
File Striping and Data Access
12.4.3.
Global Stateid Requirements
12.4.4.
The Layout Iomode
12.4.5.
Storage Device State Propagation
12.4.6.
Storage Device Component File Size
12.4.7.
Crash Recovery Considerations
12.4.8.
Security Considerations for the File Layout Type
12.4.9.
Alternate Approaches
13.
Internationalization
13.1.
Stringprep profile for the utf8str_cs type
13.2.
Stringprep profile for the utf8str_cis type
13.3.
Stringprep profile for the utf8str_mixed type
13.4.
UTF-8 Related Errors
14.
Error Values
14.1.
Error Definitions
14.2.
Operations and their valid errors
14.3.
Callback operations and their valid errors
14.4.
Errors and the operations that use them
15.
NFS version 4.1 Procedures
15.1.
Procedure 0: NULL - No Operation
15.2.
Procedure 1: COMPOUND - Compound Operations
16.
NFS version 4.1 Operations
16.1.
Operation 3: ACCESS - Check Access Rights
16.2.
Operation 4: CLOSE - Close File
16.3.
Operation 5: COMMIT - Commit Cached Data
16.4.
Operation 6: CREATE - Create a Non-Regular File Object
16.5.
Operation 7: DELEGPURGE - Purge Delegations Awaiting Recovery
16.6.
Operation 8: DELEGRETURN - Return Delegation
16.7.
Operation 9: GETATTR - Get Attributes
16.8.
Operation 10: GETFH - Get Current Filehandle
16.9.
Operation 11: LINK - Create Link to a File
16.10.
Operation 12: LOCK - Create Lock
16.11.
Operation 13: LOCKT - Test For Lock
16.12.
Operation 14: LOCKU - Unlock File
16.13.
Operation 15: LOOKUP - Lookup Filename
16.14.
Operation 16: LOOKUPP - Lookup Parent Directory
16.15.
Operation 17: NVERIFY - Verify Difference in Attributes
16.16.
Operation 18: OPEN - Open a Regular File
16.17.
Operation 19: OPENATTR - Open Named Attribute Directory
16.18.
Operation 21: OPEN_DOWNGRADE - Reduce Open File Access
16.19.
Operation 22: PUTFH - Set Current Filehandle
16.20.
Operation 23: PUTPUBFH - Set
Public Filehandle
16.21.
Operation 24: PUTROOTFH - Set Root Filehandle
16.22.
Operation 25: READ - Read from File
16.23.
Operation 26: READDIR - Read Directory
16.24.
Operation 27: READLINK - Read Symbolic Link
16.25.
Operation 28: REMOVE - Remove File System Object
16.26.
Operation 29: RENAME - Rename Directory Entry
16.27.
Operation 31: RESTOREFH - Restore Saved Filehandle
16.28.
Operation 32: SAVEFH - Save Current Filehandle
16.29.
Operation 33: SECINFO - Obtain Available Security
16.30.
Operation 34: SETATTR - Set Attributes
16.31.
Operation 37: VERIFY - Verify Same Attributes
16.32.
Operation 38: WRITE - Write to File
16.33.
Operation 40: BACKCHANNEL_CTL - Backchannel control
16.34.
Operation 41: BIND_CONN_TO_SESSION
16.35.
Operation 42: EXCHANGE_ID - Instantiate Clientid
16.36.
Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid
16.37.
Operation 44: DESTROY_SESSION - Destroy existing session
16.38.
Operation 45: FREE_STATEID - Free stateid with no locks
16.39.
Operation 46: GET_DIR_DELEGATION - Get a directory delegation
16.40.
Operation 47: GETDEVICEINFO - Get Device Information
16.41.
Operation 48: GETDEVICELIST
16.42.
Operation 49: LAYOUTCOMMIT - Commit writes made using a layout
16.43.
Operation 50: LAYOUTGET - Get Layout Information
16.44.
Operation 51: LAYOUTRETURN - Release Layout Information
16.45.
Operation 52: SECINFO_NO_NAME - Get Security on Unnamed Object
16.46.
Operation 53: SEQUENCE - Supply per-procedure sequencing and control
16.47.
Operation 54: SET_SSV
16.48.
Operation 55: TEST_STATEID - Test stateids for validity
16.49.
Operation 56: WANT_DELEGATION
16.50.
Operation 10044: ILLEGAL - Illegal operation
17.
NFS version 4.1 Callback Procedures
17.1.
Procedure 0: CB_NULL - No Operation
17.2.
Procedure 1: CB_COMPOUND - Compound Operations
18.
NFS version 4.1 Callback Operations
18.1.
Operation 3: CB_GETATTR - Get Attributes
18.2.
Operation 4: CB_RECALL - Recall an Open Delegation
18.3.
Operation 5: CB_LAYOUTRECALL
18.4.
Operation 6: CB_NOTIFY - Notify directory changes
18.5.
Operation 7: CB_PUSH_DELEG
18.6.
Operation 8: CB_RECALL_ANY - Keep any N delegations
18.7.
Operation 9: CB_RECALLABLE_OBJ_AVAIL
18.8.
Operation 10: CB_RECALL_SLOT - change flow control limits
18.9.
Operation 11: CB_SEQUENCE - Supply callback channel sequencing and control
18.10.
Operation 12: CB_WANTS_CANCELLED
18.11.
Operation 13: CB_NOTIFY_LOCK - Notify of possible lock availability
18.12.
Operation 10044: CB_ILLEGAL - Illegal Callback Operation
19.
Security Considerations
20.
IANA Considerations
20.1.
Defining new layout types
21.
References
21.1.
Normative References
21.2.
Informative References
Appendix A.
Acknowledgments
§
Authors' Addresses
§
Intellectual Property and Copyright Statements
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The NFSv4.1 protocol is a minor version of the NFSv4 protocol described in [2] (Shepler, S., Callaghan, B., Robinson, D., Thurlow, R., Beame, C., Eisler, M., and D. Noveck, “Network File System (NFS) version 4 Protocol,” April 2003.). It generally follows the guidelines for minor versioning model laid in Section 10 of RFC 3530. However, it diverges from guidelines 11 ("a client and server that supports minor version X must support minor versions 0 through X-1"), and 12 ("no features may be introduced as mandatory in a minor version"). These divergences are due to the introduction of the sessions model for managing non-idempotent operations and the RECLAIM_COMPLETE operation. These two new features are infrastructural in nature and simplify implementation of existing and other new features. Making them optional would add undue complexity to protocol definition and implementation. NFSv4.1 accordingly updates the Minor Versioning guidelines (Minor Versioning).
NFSv4.1, as a minor version, is consistent with the overall goals for NFS Version 4, but extends the protocol so as to better meet those goals, based on experiences with NFSv4.0. In addition, NFSv4.1 has adopted some additional goals, which motivate some of the major extensions in minor version 1.
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The NFS version 4 protocol is a further revision of the NFS protocol defined already by versions 2 [17] (Nowicki, B., “NFS: Network File System Protocol specification,” March 1989.)] and 3 [18] (Callaghan, B., Pawlowski, B., and P. Staubach, “NFS Version 3 Protocol Specification,” June 1995.). It retains the essential characteristics of previous versions: design for easy recovery, independent of transport protocols, operating systems and file systems, simplicity, and good performance. The NFS version 4 revision has the following goals:
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Minor version one has the following goals, within the framework established by the overall version 4 goals.
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To provide a reasonable context for the reader, the major features of NFS version 4.1 protocol will be reviewed in brief. This will be done to provide an appropriate context for both the reader who is familiar with the previous versions of the NFS protocol and the reader that is new to the NFS protocols. For the reader new to the NFS protocols, there is still a set of fundamental knowledge that is expected. The reader should be familiar with the XDR and RPC protocols as described in [3] (Eisler, M., “XDR: External Data Representation Standard,” May 2006.) and [4] (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.). A basic knowledge of file systems and distributed file systems is expected as well.
This description of version 4.1 features will not distinguish those added in minor version one from those present in the base protocol but will treat minor version 1 as a unified whole. See Section 1.6 (Differences from NFSv4.0) for a description of the differences between the two minor versions.
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As with previous versions of NFS, the External Data Representation (XDR) and Remote Procedure Call (RPC) mechanisms used for the NFS version 4.1 protocol are those defined in [3] (Eisler, M., “XDR: External Data Representation Standard,” May 2006.) and [4] (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.). To meet end-to-end security requirements, the RPCSEC_GSS framework [5] (Eisler, M., Chiu, A., and L. Ling, “RPCSEC_GSS Protocol Specification,” September 1997.) will be used to extend the basic RPC security. With the use of RPCSEC_GSS, various mechanisms can be provided to offer authentication, integrity, and privacy to the NFS version 4 protocol. Kerberos V5 will be used as described in [6] (Linn, J., “The Kerberos Version 5 GSS-API Mechanism,” June 1996.) to provide one security framework. The LIPKEY and SPKM-3 GSS-API mechanisms described in [7] (Eisler, M., “LIPKEY - A Low Infrastructure Public Key Mechanism Using SPKM,” June 2000.) will be used to provide for the use of user password and client/server public key certificates by the NFS version 4 protocol. With the use of RPCSEC_GSS, other mechanisms may also be specified and used for NFS version 4.1 security.
To enable in-band security negotiation, the NFS version 4.1 protocol has operations which provide the client a method of querying the server about its policies regarding which security mechanisms must be used for access to the server's file system resources. With this, the client can securely match the security mechanism that meets the policies specified at both the client and server.
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Unlike NFS Versions 2 and 3, which used a series of ancillary protocols (e.g. NLM, NSM, MOUNT), within all minor versions of NFS version 4 only a single RPC protocol is used to make requests of the server. Facilities that had been separate protocols, such as locking, are now integrated within a single unified protocol.
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Minor version one supports high-performance data access to a clustered server implementation by enabling a separation of metadata access and data access, with the latter done to multiple servers in parallel.
Such parallel data access is controlled by recallable objects known as "layouts", which are integrated into the protocol locking model. Clients direct requests for data access to a set of data servers specified by the layout via a data storage protocol which may be NFSv4.1 or may be another protocol.
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The general file system model used for the NFS version 4.1 protocol is the same as previous versions. The server file system is hierarchical with the regular files contained within being treated as opaque byte streams. In a slight departure, file and directory names are encoded with UTF-8 to deal with the basics of internationalization.
The NFS version 4.1 protocol does not require a separate protocol to provide for the initial mapping between path name and filehandle. All file systems exported by a server are presented as a tree so that all file systems are reachable from a special per-server global root filehandle. This allows LOOKUP operations to be used to perform functions previously provided by the MOUNT protocol. The server provides any necessary pseudo filesystems to bridge any gaps that arise due unexported gaps between exported file systems.
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As in previous versions of the NFS protocol, opaque filehandles are used to identify individual files and directories. Lookup-type and create operations are used to go from file and directory names to the filehandle which is then used to identify the object to subsequent operations.
The NFS version 4.1 protocol provides support for both persistent filehandles, guaranteed to be valid for the lifetime of the file system object designated. In addition it provides support to servers to provide filehandles with more limited validity guarantees, called volatile filehandles.
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The NFS version 4.1 protocol has a rich and extensible attribute structure. Only a small set of the defined attributes are mandatory and must be provided by all server implementations. The other attributes are known as "recommended" attributes.
One significant recommended file attribute is the Access Control List (ACL) attribute. This attribute provides for directory and file access control beyond the model used in NFS Versions 2 and 3. The ACL definition allows for specification specific sets of permissions for individual users and groups. In addition, ACL inheritance allows propagation of access permissions and restriction down a directory tree as filesystem objects are created.
One other type of attribute is the named attribute. A named attribute is an opaque byte stream that is associated with a directory or file and referred to by a string name. Named attributes are meant to be used by client applications as a method to associate application specific data with a regular file or directory.
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NFS Version 4.1 contains a number of features to allow implementation of namespaces that cross server boundaries and that allow to and facilitate a non-disruptive transfer of support for individual file systems between servers. They are all based upon attributes that allow one file system to specify alternate or new locations for that file system.
These attributes may be used together with the concept of absent file system which provide specifications for additional locations but no actual file system content. This allows a number of important facilities:
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As mentioned previously, NFS v4.1, is a single protocol which includes locking facilities. These locking facilities include support for many types of locks including a number of sorts of recallable locks. Recallable locks such as delegations allow the client to be assured that certain events will not occur so long as that lock is held. When circumstances change, the lock is recalled via a callback via a callback request. The assurances provided by delegations allow more extensive caching to be done safely when circumstances allow it.
All locks for a given client are tied together under a single client-wide lease. All requests made on sessions associated with the client renew that lease. When leases are not promptly renewed lock are subject to revocation. In the event of server reinitialization, clients have the opportunity to safely reclaim their locks within a special grace period.
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The following definitions are provided for the purpose of providing an appropriate context for the reader.
- Client
- The "client" is the entity that accesses the NFS server's resources. The client may be an application which contains the logic to access the NFS server directly. The client may also be the traditional operating system client remote file system services for a set of applications.
In the case of file locking the client is the entity that maintains a set of locks on behalf of one or more applications. This client is responsible for crash or failure recovery for those locks it manages.
Note that multiple clients may share the same transport and multiple clients may exist on the same network node.- Clientid
- A 64-bit quantity used as a unique, short-hand reference to a client supplied Verifier and ID. The server is responsible for supplying the Clientid.
- Lease
- An interval of time defined by the server for which the client is irrevocably granted a lock. At the end of a lease period the lock may be revoked if the lease has not been extended. The lock must be revoked if a conflicting lock has been granted after the lease interval.
All leases granted by a server have the same fixed interval. Note that the fixed interval was chosen to alleviate the expense a server would have in maintaining state about variable length leases across server failures.- Lock
- The term "lock" is used to refer any of record (byte- range) locks, share reservations, delegations or layouts unless specifically stated otherwise.
- Server
- The "Server" is the entity responsible for coordinating client access to a set of file systems.
- Stable Storage
- NFS version 4 servers must be able to recover without data loss from multiple power failures (including cascading power failures, that is, several power failures in quick succession), operating system failures, and hardware failure of components other than the storage medium itself (for example, disk, nonvolatile RAM).
Some examples of stable storage that are allowable for an NFS server include:
- Media commit of data, that is, the modified data has been successfully written to the disk media, for example, the disk platter.
- An immediate reply disk drive with battery-backed on- drive intermediate storage or uninterruptible power system (UPS).
- Server commit of data with battery-backed intermediate storage and recovery software.
- Cache commit with uninterruptible power system (UPS) and recovery software.
- Stateid
- A 128-bit quantity returned by a server that uniquely defines the open and locking state provided by the server for a specific open or lock owner for a specific file. meaning and are reserved values.
- Verifier
- A 64-bit quantity generated by the client that the server can use to determine if the client has restarted and lost all previous lock state.
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The following summarizes the differences between minor version one and the base protocol:
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NFS version 4.1 (NFSv4.1) relies on core infrastructure common to nearly every operation. This core infrastructure is described in the remainder of this section.
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The NFS version 4.1 (NFSv4.1) protocol is a Remote Procedure Call (RPC) application that uses RPC version 2 and the corresponding eXternal Data Representation (XDR) as defined in RFC1831 (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.) [4] and RFC4506 (Eisler, M., “XDR: External Data Representation Standard,” May 2006.) [3].
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Previous NFS versions have been thought of as having a host-based authentication model, where the NFS server authenticates the NFS client, and trust the client to authenticate all users. Actually, NFS has always depended on RPC for authentication. The first form of RPC authentication which required a host-based authentication approach. NFSv4 also depends on RPC for basic security services, and mandates RPC support for a user-based authentication model. The user-based authentication model has user principals authenticated by a server, and in turn the server authenticated by user principals. RPC provides some basic security services which are used by NFSv4.
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As described in section 7.2 "Authentication" of [4] (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.), RPC security is encapsulated in the RPC header, via a security or authentication flavor, and information specific to the specification of the security flavor. Every RPC header conveys information used to identify and authenticate a client and server. As discussed in Section 2.2.1.1.1 (RPCSEC_GSS and Security Services), some security flavors provide additional security services.
NFSv4 clients and servers MUST implement RPCSEC_GSS. (This requirement to implement is not a requirement to use.) Other flavors, such as AUTH_NONE, and AUTH_SYS, MAY be implemented as well.
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RPCSEC_GSS ([5] (Eisler, M., Chiu, A., and L. Ling, “RPCSEC_GSS Protocol Specification,” September 1997.)) uses the functionality of GSS-API RFC2743 (Linn, J., “Generic Security Service Application Program Interface Version 2, Update 1,” January 2000.) [8]. This allows for the use of various security mechanisms by the RPC layer without the additional implementation overhead of adding RPC security flavors.
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Via the GSS-API, RPCSEC_GSS can be used to identify and authenticate users on clients to servers, and servers to users. It can also perform integrity checking on the entire RPC message, including the RPC header, and the arguments or results. Finally, privacy, usually via encryption, is a service available with RPCSEC_GSS. Privacy is performed on the arguments and results. Note that if privacy is selected, integrity, authentication, and identification are enabled. If privacy is not selected, but integrity is selected, authentication and identification are enabled. If integrity and privacy are not selected, but authentication is enabled, identification is enabled. RPCSEC_GSS does not provide identification as a separate service.
Although GSS-API has an authentication service distinct from its privacy and integrity services, use GSS-API's authentication service is not used for RPCSEC_GSS's authentication service. Instead, each RPC request and response header is integrity protected with the GSS-API integrity service, and this allows RPCSEC_GSS to offer per-RPC authentication and identity. See [5] (Eisler, M., Chiu, A., and L. Ling, “RPCSEC_GSS Protocol Specification,” September 1997.) for more information.
NFSv4 client and servers MUST support RPCSEC_GSS's integrity and authentication service. NFSv4.1 servers MUST support RPCSEC_GSS's privacy service.
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RPCSEC_GSS, via GSS-API, normalizes access to mechanisms that provide security services. Therefore NFSv4 clients and servers MUST support three security mechanisms: Kerberos V5, SPKM-3, and LIPKEY.
The use of RPCSEC_GSS requires selection of: mechanism, quality of protection (QOP), and service (authentication, integrity, privacy). For the mandated security mechanisms, NFSv4 specifies that a QOP of zero (0) is used, leaving it up to the mechanism or the mechanism's configuration to use an appropriate level of protection that QOP zero maps to. Each mandated mechanism specifies minimum set of cryptographic algorithms for implementing integrity and privacy. NFSv4 clients and servers MUST be implemented on operating environments that comply with the mandatory cryptographic algorithms of each mandated mechanism.
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The Kerberos V5 GSS-API mechanism as described in RFC1964 (Linn, J., “The Kerberos Version 5 GSS-API Mechanism,” June 1996.) [6] ( [Comment.1] (need new Kerberos RFC) ) MUST be implemented with the RPCSEC_GSS services as specified in the following table:
column descriptions: 1 == number of pseudo flavor 2 == name of pseudo flavor 3 == mechanism's OID 4 == RPCSEC_GSS service 5 == NFSv4.1 clients MUST support 6 == NFSv4.1 servers MUST support 1 2 3 4 5 6 ------------------------------------------------------------------ 390003 krb5 1.2.840.113554.1.2.2 rpc_gss_svc_none yes yes 390004 krb5i 1.2.840.113554.1.2.2 rpc_gss_svc_integrity yes yes 390005 krb5p 1.2.840.113554.1.2.2 rpc_gss_svc_privacy no yes
Note that the number and name of the pseudo flavor is presented here as a mapping aid to the implementor. Because the NFSv4 protocol includes a method to negotiate security and it understands the GSS-API mechanism, the pseudo flavor is not needed. The pseudo flavor is needed for the NFS version 3 since the security negotiation is done via the MOUNT protocol as described in [19] (Eisler, M., “NFS Version 2 and Version 3 Security Issues and the NFS Protocol's Use of RPCSEC_GSS and Kerberos V5,” June 1999.).
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The LIPKEY V5 GSS-API mechanism as described in [7] (Eisler, M., “LIPKEY - A Low Infrastructure Public Key Mechanism Using SPKM,” June 2000.) MUST be implemented with the RPCSEC_GSS services as specified in the following table:
1 2 3 4 5 6 ------------------------------------------------------------------ 390006 lipkey 1.3.6.1.5.5.9 rpc_gss_svc_none yes yes 390007 lipkey-i 1.3.6.1.5.5.9 rpc_gss_svc_integrity yes yes 390008 lipkey-p 1.3.6.1.5.5.9 rpc_gss_svc_privacy no yes
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The SPKM-3 GSS-API mechanism as described in [7] (Eisler, M., “LIPKEY - A Low Infrastructure Public Key Mechanism Using SPKM,” June 2000.) MUST be implemented with the RPCSEC_GSS services as specified in the following table:
1 2 3 4 5 6 ------------------------------------------------------------------ 390009 spkm3 1.3.6.1.5.5.1.3 rpc_gss_svc_none yes yes 390010 spkm3i 1.3.6.1.5.5.1.3 rpc_gss_svc_integrity yes yes 390011 spkm3p 1.3.6.1.5.5.1.3 rpc_gss_svc_privacy no yes
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Regardless of what security mechanism under RPCSEC_GSS is being used, the NFS server, MUST identify itself in GSS-API via a GSS_C_NT_HOSTBASED_SERVICE name type. GSS_C_NT_HOSTBASED_SERVICE names are of the form:
service@hostname
For NFS, the "service" element is
nfs
Implementations of security mechanisms will convert nfs@hostname to various different forms. For Kerberos V5, LIPKEY, and SPKM-3, the following form is RECOMMENDED:
nfs/hostname
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A significant departure from the versions of the NFS protocol before version 4 is the introduction of the COMPOUND procedure. For the NFSv4 protocol, in all minor versions, there are exactly two RPC procedures, NULL and COMPOUND. The COMPOUND procedure is defined as a series of individual operations and these operations perform the sorts of functions performed by traditional NFS procedures.
The operations combined within a COMPOUND request are evaluated in order by the server, without any atomicity guarantees. A limited set of facilities exist to pass results from one operation to another. Once an operation returns a failing result, the evaluation ends and the results of all evaluated operations are returned to the client.
With the use of the COMPOUND procedure, the client is able to build simple or complex requests. These COMPOUND requests allow for a reduction in the number of RPCs needed for logical file system operations. For example, multi-component lookup requests can be constructed by combining multiple LOOKUP operations. Those can be further combined with operations such as GETATTR, READDIR, or OPEN plus READ to do more complicated sets of operation without incurring additional latency.
NFSv4 also contains a considerable set of callback operations in which the server makes an RPC directed at the client. Callback RPC's have a similar structure to that of the normal server requests. For the NFS version 4 protocol callbacks in all minor versions, there are two RPC procedures, NULL and CB_COMPOUND. The CB_COMPOUND procedure is defined in analogous fashion to that of COMPOUND with its own set of callback operations.
Addition of new server and callback operation within the COMPOUND and CB_COMPOUND request framework provide means of extending the protocol in subsequent minor versions.
Except for a small number of operations needed for session creation, server requests and callback requests are performed within the context of a session. Sessions provide a client context for every request and support robust replay protection for non-idempotent requests.
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For each operation that obtains or depends on locking state, the specific client must be determinable by the server. In NFSv4, each distinct client instance is represented by a clientid, which is a 64-bit identifier that identifies a specific client at a given time and which is changed whenever the client or the server re-initializes. Clientid's are used to support lock identification and crash recovery.
In NFSv4.1, the clientid associated with each operation is derived from the session (see Section 2.9 (Session)) on which the operation is issued. Each session is associated with a specific clientid at session creation and that clientid then becomes the clientid associated with all requests issued using it. Therefore, unlike NFSv4.0, no NFSv4.1 operation is possible until a clientid is established.
A sequence of an EXCHANGE_ID operation followed by a CREATE_SESSION operation using that clientid is required to establish the identification on the server. Establishment of identification by a new incarnation of the client also has the effect of immediately releasing any locking state that a previous incarnation of that same client might have had on the server. Such released state would include all lock, share reservation, and, where the server is not supporting the CLAIM_DELEGATE_PREV claim type, all delegation state associated with same client with the same identity. For discussion of delegation state recovery, see Section 9.2.1 (Delegation Recovery).
Releasing such state requires that the server be able to determine that one client instance is the successor of another. Where this cannot be done, for any of a number of reasons, the locking state will remain for a time subject to lease expiration (see Section 8.5 (Lease Renewal)) and the new client will need to wait for such state to be removed, if it makes conflicting lock requests.
Client identification is encapsulated in the following structure:
struct client_owner4 {
verifier4 co_verifier;
opaque co_ownerid<NFS4_OPAQUE_LIMIT>;
};
The first field, co_verifier, is a client incarnation verifier that is used to detect client reboots. Only if the co_verifier is different from that the server had previously recorded for the client (as identified by the second field of the structure, co_ownerid) does the server start the process of canceling the client's leased state.
The second field, co_ownerid is a variable length string that uniquely defines the client so that subsequent instances of the same client bear the same co_ownerid with a different verifier.
There are several considerations for how the client generates the co_ownerid string:
Given the above considerations, an example of a well generated co_ownerid string is one that includes:
As a security measure, the server MUST NOT cancel a client's leased state if the principal established the state for a given co_ownerid string is not the same as the principal issuing the EXCHANGE_ID.
A server may compare an client_owner4 in a EXCHANGE_ID with an nfs_client_id4 established using SETCLIENTID using NFSv4 minor version 0, so that an NFSv4.1 client is not forced to delay until lease expiration for locking state established by the earlier client using minor version 0. This requires the client_owner4 be constructed the same way as the nfs_client_id4. If the latter's contents included the server's network address, and the NFSv4.1 client does not wish to use a clientid that prevents trunking, it should issue two EXCHANGE_ID operations. The first EXCHANGE_ID will have a client_owner4 equal to the nfs_client_id4. This will clear the state created by the NFSv4.0 client. The second EXCHANGE_ID will not have the server's network address. The state created for the second EXCHANGE_ID will not have to wait for lease expiration, because there will be no state to expire.
Once a EXCHANGE_ID has been done, and the resulting clientid established as associated with a session, all requests made on that session implicitly identify that clientid, which in turn designates the client specified using the long-form client_owner4 structure. The shorthand client identifier (a clientid) is assigned by the server and should be chosen so that it will not conflict with a clientid previously assigned by the server. This applies across server restarts or reboots.
In the event of a server restart, a client will find out that its current clientid is no longer valid when receives a NFS4ERR_STALE_CLIENTID error. The precise circumstances depend of the characteristics of the sessions involved, specifically whether the session is persistent (see Section 2.9.4.5 (Persistence)).
When a session is not persistent, the client will need to create a new session. When the existing clientid is presented to a server as part of creating a session and that clientid is not recognized, as would happen after a server reboot, the server will reject the request with the error NFS4ERR_STALE_CLIENTID. When this happens, the client must obtain a new clientid by use of the EXCHANGE_ID operation and then use that clientid as the basis of the basis of a new session and then proceed to any other necessary recovery for the server reboot case (See Section 8.6.2 (Server Failure and Recovery)).
In the case of the session being persistent, the client will re-establish communication using the existing session after the reboot. This session will be associated with a stale clientid and the client will receive an indication of that fact in the sr_status field returned by the SEQUENCE operation (see Section 2.9.2.1 (SEQUENCE and CB_SEQUENCE)). The client can then use the existing session to do whatever operations are necessary to determine the status of requests outstanding at the time of reboot, while avoiding issuing new requests, particularly any involving locking on that session. Such requests would fail with NFS4ERR_STALE_CLIENTID error or an NFS4ERR_STALE_STATEID error, if attempted. In any case, the client would create a new clientid using EXCHANGE_ID, create a new session based on that clientid, and proceed to other necessary recovery for the server reboot case.
See the detailed descriptions of EXCHANGE_ID (Section 16.35 (Operation 42: EXCHANGE_ID - Instantiate Clientid) and CREATE_SESSION (Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid)) for a complete specification of these operations.
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If the server determines that the client holds no associated state for its clientid, the server may choose to release the clientid. The server may make this choice for an inactive client so that resources are not consumed by those intermittently active clients. If the client contacts the server after this release, the server must ensure the client receives the appropriate error so that it will use the EXCHANGE_ID/CREATE_SESSION sequence to establish a new identity. It should be clear that the server must be very hesitant to release a clientid since the resulting work on the client to recover from such an event will be the same burden as if the server had failed and restarted. Typically a server would not release a clientid unless there had been no activity from that client for many minutes. Note that "associated state" includes sessions. As long as there are sessions, the server MUST not release the clientid. See Section 2.9.8.1.4 (Loss of Session) for discussion on releasing inactive sessions.
Note that if the id string in a EXCHANGE_ID request is properly constructed, and if the client takes care to use the same principal for each successive use of EXCHANGE_ID, then, barring an active denial of service attack, NFS4ERR_CLID_INUSE should never be returned.
However, client bugs, server bugs, or perhaps a deliberate change of the principal owner of the id string (such as the case of a client that changes security flavors, and under the new flavor, there is no mapping to the previous owner) will in rare cases result in NFS4ERR_CLID_INUSE.
In that event, when the server gets a EXCHANGE_ID for a client id that currently has no state, or it has state, but the lease has expired, rather than returning NFS4ERR_CLID_INUSE, the server MUST allow the EXCHANGE_ID, and confirm the new clientid if followed by the appropriate CREATE_SESSION.
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With the NFS version 4 server potentially offering multiple security mechanisms, the client needs a method to determine or negotiate which mechanism is to be used for its communication with the server. The NFS server may have multiple points within its file system namespace that are available for use by NFS clients. These points can be considered security policy boundaries, and in some NFS implementations are tied to NFS export points. In turn the NFS server may be configured such that each of these security policy boundaries may have different or multiple security mechanisms in use.
The security negotiation between client and server must be done with a secure channel to eliminate the possibility of a third party intercepting the negotiation sequence and forcing the client and server to choose a lower level of security than required or desired. See Section 19 (Security Considerations) for further discussion.
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An NFS server can assign one or more "security tuples" to each security policy boundary in its namespace. Each security tuple consists of a security flavor (see Section 2.2.1.1 (RPC Security Flavors)), and if the flavor is RPCSEC_GSS, a GSS-API mechanism OID, a GSS-API quality of protection, and an RPCSEC_GSS service.
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The SECINFO and SECINFO_NO_NAME operations allow the client to determine, on a per filehandle basis, what security tuple is to be used for server access. In general, the client will not have to use either operation except during initial communication with the server or when the client crosses security policy boundaries at the server. It is possible that the server's policies change during the client's interaction therefore forcing the client to negotiate a new security tuple.
Where the use of different security tuples would affect the type of access that would be allowed if a request was issued over the same connection used for the SECINFO or SECINFO_NO_NAME operation (e.g. read-only vs. read-write) access, security tuples that allow greater access should be presented first. Where the general level of access is the same and different security flavors limit the range of principals whose privileges are recognized (e.g. allowing or disallowing root access), flavors supporting the greatest range of principals should be listed first.
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Based on the assumption that each NFS version 4 client and server must support a minimum set of security (i.e., LIPKEY, SPKM-3, and Kerberos-V5 all under RPCSEC_GSS), the NFS client will initiate file access to the server with one of the minimal security tuples. During communication with the server, the client may receive an NFS error of NFS4ERR_WRONGSEC. This error allows the server to notify the client that the security tuple currently being used is contravenes the server's security policy. The client is then responsible for determining (see Section 2.5.3.1 (Using NFS4ERR_WRONGSEC, SECINFO, and SECINFO_NO_NAME)) what security tuples are available at the server and choose one which is appropriate for the client.
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This section explains of the mechanics of NFSv4.1 security negotiation. Unless noted otherwise, for any mention of PUTFH in this section, the reader should interpret it as applying to PUTROOTFH and PUTPUBFH in addition to PUTFH.
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This situation also applies to a put filehandle operation followed by an OPEN operation that specifies a component name.
In this situation, the client is potentially crossing a security policy boundary, and the set of security tuples the parent directory supports differ from those of the child. The server implementation may decide whether to impose any restrictions on security policy administration. There are at least three approaches (sec_policy_child is the tuple set of the child export, sec_policy_parent is that of the parent).
- a)
- sec_policy_child <= sec_policy_parent (<= for subset). This means that the set of security tuples specified on the security policy of a child directory is always a subset of that of its parent directory.
- b)
- sec_policy_child ^ sec_policy_parent != {} (^ for intersection, {} for the empty set). This means that the security tuples specified on the security policy of a child directory always has a non empty intersection with that of the parent.
- c)
- sec_policy_child ^ sec_policy_parent == {}. This means that the set of tuples specified on the security policy of a child directory may not intersect with that of the parent. In other words, there are no restrictions on how the system administrator may set.
For a server to support approach (b) (when client chooses a flavor that is not a member of sec_policy_parent) and (c), PUTFH must NOT return NFS4ERR_WRONGSEC in case of security mismatch. Instead, it should be returned from the LOOKUP (or OPEN by component name) that follows.
Since the above guideline does not contradict approach (a), it should be followed in general. Even if approach (a) is implemented, it is possible for the security tuple used to be acceptable for the target of LOOKUP but not for the filehandles used in PUTFH. The PUTFH could really be a PUTROOTFH or PUTPUBFH, where the client does not know the security tuples for the root or public filehandle. Or the security policy for the filehandle used by PUTFH could have changed since the time the filehandle was obtained.
Therefore, an NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC in response to PUTFH, PUTROOTFH, or PUTPUBFH if the operation is immediately followed by a LOOKUP or an OPEN by component name.
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Since SECINFO only works its way down, there is no way LOOKUPP can return NFS4ERR_WRONGSEC without SECINFO_NO_NAME. SECINFO_NO_NAME solves this issue because via style "parent", it works in the opposite direction as SECINFO. As with Section 2.5.3.1.1 (PUTFH + LOOKUP (or OPEN by Name)), PUTFH must not return NFS4ERR_WRONGSEC whenever it is followed by LOOKUPP. If the server does not support SECINFO_NO_NAME, the client's only recourse is to issue the PUTFH, LOOKUPP, GETFH sequence of operations with every security tuple it supports.
Regardless whether SECINFO_NO_NAME is supported, an NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC in response to PUTFH, PUTROOTFH, or PUTPUBFH if the operation is immediately followed by a LOOKUPP.
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A security sensitive client is allowed to choose a strong security tuple when querying a server to determine a file object's permitted security tuples. The security tuple chosen by the client does not have to be included in the tuple list of the security policy of the either parent directory indicated in PUTFH, or the child file object indicated in SECINFO (or any parent directory indicated in SECINFO_NO_NAME). Of course the server has to be configured for whatever security tuple the client selects, otherwise the request will fail at RPC layer with an appropriate authentication error.
In theory, there is no connection between the security flavor used by SECINFO or SECINFO_NO_NAME and those supported by the security policy. But in practice, the client may start looking for strong flavors from those supported by the security policy, followed by those in the mandatory set.
The NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC to PUTFH whenever it is immediately followed by SECINFO or SECINFO_NO_NAME. The NFSv4.1 server MUST NOT return NFS4ERR_WRONGSEC from SECINFO or SECINFO_NO_NAME.
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This is a nonsensical situation, because the first put filehandle operation is wasted. The NFSv4.1 server MAY return NFS4ERR_WRONGSEC to the first PUTFH, or it MAY NOT. If it does not, it then processes the subsequent PUTFH and any operation that follows it according to the rules listed in Section 2.5.3.1 (Using NFS4ERR_WRONGSEC, SECINFO, and SECINFO_NO_NAME).
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This too is nonsensical because the PUTFH is wasted. The NFSv4.1 server MAY or MAY NOT return NFS4ERR_WRONGSEC.
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"Anything Else" includes OPEN by filehandle.
The security policy enforcement applies to the filehandle specified in PUTFH. Therefore PUTFH must return NFS4ERR_WRONGSEC in case of security tuple on the part of the mismatch. This avoids the complexity adding NFS4ERR_WRONGSEC as an allowable error to every other operation.
PUTFH + SECINFO_NO_NAME (style "current_fh") is an efficient way for the client to recover from NFS4ERR_WRONGSEC.
The NFSv4.1 server, MUST not return NFS4ERR_WRONGSEC to any operation other than LOOKUP, LOOKUPP, and OPEN (by component name).
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To address the requirement of an NFS protocol that can evolve as the need arises, the NFS version 4 protocol contains the rules and framework to allow for future minor changes or versioning.
The base assumption with respect to minor versioning is that any future accepted minor version must follow the IETF process and be documented in a standards track RFC. Therefore, each minor version number will correspond to an RFC. Minor version zero of the NFS version 4 protocol is represented by [2] (Shepler, S., Callaghan, B., Robinson, D., Thurlow, R., Beame, C., Eisler, M., and D. Noveck, “Network File System (NFS) version 4 Protocol,” April 2003.), and minor version one is represented by this document [Comment.2] (change "document" to "RFC" when we publish) . The COMPOUND and CB_COMPOUND procedures support the encoding of the minor version being requested by the client.
The following items represent the basic rules for the development of minor versions. Note that a future minor version may decide to modify or add to the following rules as part of the minor version definition.
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As described in Section 2.2.1.1.1.1 (Identification, Authentication, Integrity, Privacy), NFSv4 relies on RPC for identification, authentication, integrity, and privacy. NFSv4 itself provides additional security services as described in the next several subsections.
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Authorization to access a file object via an NFSv4 operation is ultimately determined by the NFSv4 server. A client can predetermine its access to a file object via the OPEN (Section 16.16 (Operation 18: OPEN - Open a Regular File)) and the ACCESS (Section 16.1 (Operation 3: ACCESS - Check Access Rights)) operations.
Principals with appropriate access rights can modify the authorization on a file object via the SETATTR (Section 16.30 (Operation 34: SETATTR - Set Attributes)) operation. Four attributes that affect access rights are: mode, owner, owner_group, and acl. See Section 5 (File Attributes).
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NFSv4 provides auditing on a per file object basis, via the ACL attribute as described in Section 6 (Access Control Lists). It is outside the scope of this specification to specify audit log formats or management policies.
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NFSv4 provides alarm control on a per file object basis, via the ACL attribute as described in Section 6 (Access Control Lists). Alarms may serve as the basis for instrusion detection. It is outside the scope of this specification to specify heuristics for detecting intrusion via alarms.
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NFSv4 works over RDMA and non-RDMA_based transports with the following attributes:
Where an NFS version 4 implementation supports operation over the IP network protocol, any transport used between NFS and IP MUST be among the IETF-approved congestion control transport protocols. At the time this document was written, the only two transports that had the above attributes were TCP and SCTP. To enhance the possibilities for interoperability, an NFS version 4 implementation MUST support operation over the TCP transport protocol.
Even if NFS version 4 is used over a non-IP network protocol, it is RECOMMENDED that the transport support congestion control.
Note that it is permissible for connectionless transports to be used under NFSv4.1, however reliable and in-order delivery of data is still required. NFSv4.1 assumes that a client transport address and server transport address used to send data over a transport together constitute a connection, even if the underlying transport eschews the concept of a connection.
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If a connection-oriented transport (e.g. TCP) is used the client and server SHOULD use long lived connections for at least three reasons:
In order to reduce congestion, if a connection-oriented transport is used, and the request is not the NULL procedure,
When using RDMA transports there are other reasons not tolerating retries over the same connection:
In addition, the sender of an NFSv4.1 request is not allowed to stop waiting for a reply, as described in Section 2.9.4.2 (Retry and Replay).
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Historically, NFS version 2 and version 3 servers have resided on port 2049. The registered port 2049 RFC3232 (Reynolds, J., “Assigned Numbers: RFC 1700 is Replaced by an On-line Database,” January 2002.) [21] for the NFS protocol should be the default configuration. NFSv4 clients SHOULD NOT use the RPC binding protocols as described in RFC1833 (Srinivasan, R., “Binding Protocols for ONC RPC Version 2,” August 1995.) [22].
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Previous versions and minor versions of NFS have suffered from the following:
Through the introduction of a session, NFSv4.1 addresses the above shortfalls with practical solutions:
A session is a dynamically created, long-lived server object created by a client, used over time from one or more transport connections. Its function is to maintain the server's state relative to the connection(s) belonging to a client instance. This state is entirely independent of the connection itself, and indeed the state exists whether the connection exists or not (though locks, delegations, etc. and generally expire in the extended absence of an open connection). The session in effect becomes the object representing an active client on a set of zero or more connections.
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Sessions are part of NFSv4.1 and not NFSv4.0. Normally, a major infrastructure change like sessions would require a new major version number to an RPC program like NFS. However, because NFSv4 encapsulates its functionality in a single procedure, COMPOUND, and because COMPOUND can support an arbitrary number of operations, sessions are almost trivially added. COMPOUND includes a minor version number field, and for NFSv4.1 this minor version is set to 1. When the NFSv4 server processes a COMPOUND with the minor version set to 1, it expects a different set of operations than it does for NFSv4.0. One operation it expects is the SEQUENCE operation, which is required for every COMPOUND that operates over an established session.
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In NFSv4.1, when the SEQUENCE operation is present, it is always the first operation in the COMPOUND procedure. The primary purpose of SEQUENCE is to carry the session identifier. The session identifier associates all other operations in the COMPOUND procedure with a particular session. SEQUENCE also contains required information for maintaining EOS (see Section 2.9.4 (Exactly Once Semantics)). Session-enabled NFSv4.1 COMPOUND requests thus have the form:
+-----+--------------+-----------+------------+-----------+----
| tag | minorversion | numops |SEQUENCE op | op + args | ...
| | (== 1) | (limited) | + args | |
+-----+--------------+-----------+------------+-----------+----
and the reply's structure is:
+------------+-----+--------+-------------------------------+--//
|last status | tag | numres |status + SEQUENCE op + results | //
+------------+-----+--------+-------------------------------+--//
//-----------------------+----
// status + op + results | ...
//-----------------------+----
A CB_COMPOUND procedure request and reply has a similar form, but instead of a SEQUENCE operation, there is a CB_SEQUENCE operation, and there is an additional field called "callback_ident", which is superfluous in NFSv4.1. CB_SEQUENCE has the same information as SEQUENCE, but includes other information needed to solve callback races (Section 2.9.4.3 (Resolving server callback races with sessions)).
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Sessions are subordinate to the clientid (Section 2.4 (Client Identifiers)). Each clientid can have zero or more active sessions. A clientid, and a session bound to it are required to do anything useful in NFSv4.1. Each time a session is used, the state leased to it associated clientid is automatically renewed.
State such as share reservations, locks, delegations, and layouts (Section 1.4.4 (Locking Facilities)) is tied to the clientid, not the sessions of the clientid. Successive state changing operations from a given state owner can go over different sessions, as long each session is associated with the same clientid. Callbacks can arrive over a different session than the session that sent the operation the acquired the state that the callback is for. For example, if session A is used to acquire a delegation, a request to recall the delegation can arrive over session B.
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Each session has one or two channels: the "operation" or "fore" channel used for ordinary requests from client to server, and the "back" channel, used for callback requests from server to client. The session allocates resources for each channel, including separate reply caches (see Section 2.9.4.1 (Slot Identifiers and Reply Cache) These resources are for the most part specified at time the session is created.
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The operation channel carries COMPOUND requests and responses. A session always has an operation channel.
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The backchannel carries CB_COMPOUND requests and responses. Whether there is a backchannel or not is a decision of the client; NFSv4.1 servers MUST support backchannels.
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Because there are at most two channels per session, and because each channel has a distinct purpose, channels are not assigned identifiers. The operation and backchannel are implicitly created and associated when the session is created.
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Each channel is associated with zero or more transport connections. A connection can be bound to one channel or both channels of a session; the client and server negotiate whether a connection will carry traffic for one channel or both channels via the CREATE_SESSION (Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid)) and the BIND_CONN_TO_SESSION (Section 16.34 (Operation 41: BIND_CONN_TO_SESSION)) operations. When a session is created via CREATE_SESSION, it is automatically bound to the operation channel, and optionally the backchannel. If the client does not specify connecting binding enforcement when the session is created, then additional connections are automatically bound to the operation channel when the are used with a SEQUENCE operation that has the session's sessionid.
A connection MAY be bound to the channels of other sessions. The client decides, and the NFSv4.1 server MUST allow it. A connection MAY be bound to the channels of other sessions of other clientids. Again, the client decides, and the server MUST allow it.
It is permissible for connections of multiple types to be bound to the same channel. For example a TCP and RDMA connection can be bound to the operation channel. In the event an RDMA and non-RDMA connection are bound to the same channel, the maximum number of slots must be at least one more than the total number of credits. This way if all RDMA credits are use, the non-RDMA connection can have at least one outstanding request.
It is permissible for a connection of one type to be bound to the operation channel, and another type bound to the backchannel.
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The eir_server_owner results from EXCHANGE_ID give a client a hint that the server it is connected to may be the same as the server it is connected to via another connection. When two connections have the same eir_server_owner.so_major_id, the client treats the connections as connected to the same server (even if the destination network addresses are different) and uses a common clientid to identify itself. The eir_server_owner.so_minor_id field allows the server to control binding of connections to sessions. When two connections have a matching so_major_id and so_minor_id, the client may bind both connections to a common session; this is session trunking. When two connections have a matching so_major_id, but different so_minor_id, the client will need to create a new session for the clientid in order to use the connection; this is clientid trunking. In either session or clientid trunking, the bandwidth capacity can scale with the number of connections.
Just because two servers over two connections claim matching or partially matching server_owner4 values does not the client should or must trust the servers' claims. The client may verify these claims before trunking traffic.
For session trunking, clients and servers can reliably verify if connections between different network paths are in fact bound to the same NFSv4.1 server and usable on the same session. The SET_SSV (Section 16.47 (Operation 54: SET_SSV)) operation allows a client and server to establish a unique, shared key value (the SSV). When a new connection is bound to the session (via the BIND_CONN_TO_SESSION operation, see Section 16.34 (Operation 41: BIND_CONN_TO_SESSION)), the client offers a digest that based on the SSV. If the client mistakenly tries to bind a connection to a session of a wrong server, the server will either reject the attempt because it is not aware of the session identifier of the BIND_CONN_TO_SESSION arguments, or it will reject the attempt because the digest for the SSV does not match what the server expects. Even if the server mistakenly or maliciously accepts the connection bind attempt, the digest it computes in the response will not be verified by the client, the client will know it cannot use the connection for trunking the specified channel.
In the case of clientid trunking, the client can use RPCSEC_GSS to verify that each connection is aimed at the same server. When the client invokes EXCHANGE_ID, it should use RPCSEC_GSS. If each RPCSEC_GSS context over each connection has the same server principal, then the servers at the end of each connection are the same.
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Via the session, NFSv4.1 offers exactly once semantics (EOS) for requests sent over a channel. EOS is supported on both the operation and back channels.
Each COMPOUND or CB_COMPOUND request that is issued with a leading SEQUENCE or CB_SEQUENCE operation MUST be executed by the receiver exactly once. This requirement is regardless whether the request is issued with reply caching specified (see Section 2.9.4.1.2 (Optional Reply Caching)). The requirement holds even if the requester is issuing the request over a session created between a pNFS data client and pNFS data server. The rationale for this requirement is understood by categorizing requests into three classifications:
An example of a non-idempotent request is RENAME. If is obvious that if a replier executes the same RENAME request twice, and the first execution succeeds, the re-execution will fail. If the replier returns the result from the re-execution, this result is incorrect. Therefore, EOS is required for nonidempotent requests.
An example of an idempotent modifying request is a COMPOUND request containing a WRITE operation. Repeated execution of the same WRITE has the same effect as execution of that write once. Nevertheless, putting enforcing EOS for WRITEs and other idempotent modifying requests is necessary to avoid data corruption.
Suppose a client issues WRITEs A, B, C to a noncompliant server that does not enforce EOS, and receives no response, perhaps due to a network partition. The client reconnects to the server and re-issues all three WRITEs. Now, the server has outstanding two instances of each of A, B, and C. The server can be in a situation in which it executes and replies to the retries of A, B, and C while the first A, B, and C are still waiting around in the server's I/O system for some resource. Upon receiving the replies to the second attempts of WRITEs A, B, and C, the client believes its writes are done so it is free to do issue WRITE D which overlaps the range of one or more of A, B, C. If any of A, B, or C are subsequently are executed for the second time, then what has been written by D can be overwritten and thus corrupted.
Note that it is not required the server cache the reply to the modifying operation to avoid data corruption (but if the client specified the reply to be cached, the server must cache it).
An example of an idempotent non-modifying request is a COMPOUND containing SEQUENCE, PUTFH, READLINK and nothing else. The re-execution of a such a request will not cause data corruption, or produce an incorrect result. Nonetheless, for simplicity, the replier MUST enforce EOS for such requests.
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The RPC layer provides a transaction ID (xid), which, while required to be unique, is not especially convenient for tracking requests. The xid is only meaningful to the requester it cannot be interpreted at the replier except to test for equality with previously issued requests. Because RPC operations may be completed by the replier in any order, many transaction IDs may be outstanding at any time. The requester may therefore perform a computationally expensive lookup operation in the process of demultiplexing each reply.
In the NFSv4.1, there is a limit to the number of active requests. This immediately enables a computationally efficient index for each request which is designated as a Slot Identifier, or slotid.
When the requester issues a new request, it selects a slotid in the range 0..N-1, where N is the replier's current "totalrequests" limit granted to the requester on the session over which the request is to be issued. The slotid must be unused by any of the requests which the requester has already active on the session. "Unused" here means the requester has no outstanding request for that slotid. Because the slot id is always an integer in the range 0..N-1, requester implementations can use the slotid from a replier response to efficiently match responses with outstanding requests, such as, for example, by using the slotid to index into a outstanding request array. This can be used to avoid expensive hashing and lookup functions in the performance-critical receive path.
The sequenceid, which accompanies the slotid in each request, is important for an important check at the server: it must be able to be determined efficiently whether a request using a certain slotid is a retransmit or a new, never-before-seen request. It is not feasible for the client to assert that it is retransmitting to implement this, because for any given request the client cannot know the server has seen it unless the server actually replies. Of course, if the client has seen the server's reply, the client would not retransmit.
The sequenceid MUST increase monotonically for each new transmit of a given slotid, and MUST remain unchanged for any retransmission. The server must in turn compare each newly received request's sequenceid with the last one previously received for that slotid, to see if the new request is:
Unlike the XID, the slotid is always within a specific range; this has two implications. The first implication is that for a given session, the replier need only cache the results of a limited number of COMPOUND requests. The second implication derives from the first, which is unlike XID-indexed reply caches (also know as duplicate request caches - DRCs), the slotid-based reply cache cannot be overflowed. Through use of the sequenceid to identify retransmitted requests, the replier does not need to actually cache the request itself, reducing the storage requirements of the reply cache further. These new facilities makes it practical to maintain all the required entries for an effective reply cache.
The slotid and sequenceid therefore take over the traditional role of the XID and port number in the replier reply cache implementation, and the session replaces the IP address. This approach is considerably more portable and completely robust - it is not subject to the frequent reassignment of ports as clients reconnect over IP networks. In addition, the RPC XID is not used in the reply cache, enhancing robustness of the cache in the face of any rapid reuse of XIDs by the client. [Comment.3] (We need to discuss the requirements of the client for changing the XID.) .
It is required to encode the slotid information into each request in a way that does not violate the minor versioning rules of the NFSv4.0 specification. This is accomplished here by encoding it in the SEQUENCE operation within each NFSv4.1 COMPOUND and CB_COMPOUND procedure. The operation easily piggybacks within existing messages. [Comment.4] (Need a better term than piggyback)
In general, the receipt of a new sequenced request arriving on any valid slot is an indication that the previous reply cache contents of that slot may be discarded. In order to further assist the replier in slot management, the requester is required to use the lowest available slot when issuing a new request. In this way, the replier may be able to retire additional entries.
However, in the case where the replier is actively adjusting its granted maximum request count to the requester, it may not be able to use receipt of the slotid to retire cache entries. The slotid used in an incoming request may not reflect the server's current idea of the requester's session limit, because the request may have been sent from the requester before the update was received. Therefore, in the downward adjustment case, the replier may have to retain a number of reply cache entries at least as large as the old value, until operation sequencing rules allow it to infer that the requester has seen its reply.
The SEQUENCE (and CB_SEQUENCE) operation also carries a "maxslot" value which carries additional client slot usage information. The requester must always provide its highest-numbered outstanding slot value in the maxslot argument, and the replier may reply with a new recognized value. The requester should in all cases provide the most conservative value possible, although it can be increased somewhat above the actual instantaneous usage to maintain some minimum or optimal level. This provides a way for the requester to yield unused request slots back to the replier, which in turn can use the information to reallocate resources. Obviously, maxslot can never be zero, or the session would deadlock.
The replier also provides a target maxslot value to the requester, which is an indication to the requester of the maxslot the replier wishes the requester to be using. This permits the server to withdraw (or add) resources from a requester that has been found to not be using them, in order to more fairly share resources among a varying level of demand from other requesters. The requester must always comply with the replier's value updates, since they indicate newly established hard limits on the requester's access to session resources. However, because of request pipelining, the requester may have active requests in flight reflecting prior values, therefore the replier must not immediately require the requester to comply.
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Any time SEQUENCE or CB_SEQUENCE return an error, the sequenceid of the slot MUST NOT change. The replier MUST NOT modify the reply cache entry for the slot whenever an error is returned from SEQUENCE or CB_SEQUENCE.
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On a per-request basis the requester can choose to direct the replier to cache the reply to all operations after the first operation (SEQUENCE or CB_SEQUENCE) via the sa_cachethis or csa_cachethis fields of the arguments to SEQUENCE or CB_SEQUENCE. The reason it would not direct the replier to cache the entire reply is that the request is composed of all idempotent operations [20] (Juszczak, C., “Improving the Performance and Correctness of an NFS Server,” June 1990.). Caching the reply may offer little benefit, and if the reply is too large (see Section 2.9.4.4 (COMPOUND and CB_COMPOUND Construction Issues), it may not be cacheable anyway.
Whether the requester requests the reply to be cached or not has no effect on the slot processing. If the results of SEQUENCE or CB_SEQUENCE are NFS4_OK, then the slot's sequenceid MUST be incremented by one. If a requester does not direct the replier to cache, the reply, the replier MUST do one of following:
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Multiple connections can be bound to a session's channel, hence the connections share the same table of slotids. For connections over non-RDMA transports like TCP, there are no particular considerations. Considerations for multiple RDMA connections sharing a slot table are discussed in Section 2.9.5.1 (RDMA Connection Resources). [Comment.5] (Also need to discuss when RDMA and non-RDMA share a slot table.)
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A client MUST NOT retry a request, unless the connection it used to send the request disconnects. The client can then reconnect and resend the request, or it can resend the request over a different connection. In the case of the server resending over the backchannel, it cannot reconnect, and either resends the request over another connection that the client has bound to the backchannel, or if there is no other backchannel connection, waits for the client to bind a connection to the backchannel.
A client MUST wait for a reply to a request before using the slot for another request. If it does not wait for a reply, then the client does not know what sequenceid to use for the slot on its next request. For example, suppose a client sends a request with sequenceid 1, and does not wait for the response. The next time it uses the slot, it sends the new request with sequenceid 2. If the server has not seen the request with sequenceid 1, then the server is expecting sequenceid 2, and rejects the client's new request with NFS4ERR_SEQ_MISORDERED (as the result from SEQUENCE or CB_SEQUENCE).
RDMA fabrics do not guarantee that the memory handles (Steering Tags) within each RDMA three-tuple are valid on a scope [Comment.6] (What is a three-tuple?) outside that of a single connection. Therefore, handles used by the direct operations become invalid after connection loss. The server must ensure that any RDMA operations which must be replayed from the reply cache use the newly provided handle(s) from the most recent request.
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It is possible for server callbacks to arrive at the client before the reply from related forward channel operations. For example, a client may have been granted a delegation to a file it has opened, but the reply to the OPEN (informing the client of the granting of the delegation) may be delayed in the network. If a conflicting operation arrives at the server, it will recall the delegation using the callback channel, which may be on a different transport connection, perhaps even a different network. In NFSv4.0, if the callback request arrives before the related reply, the client may reply to the server with an error.
The presence of a session between client and server alleviates this issue. When a session is in place, each client request is uniquely identified by its { slotid, sequenceid } pair. By the rules under which slot entries (reply cache entries) are retired, the server has knowledge whether the client has "seen" each of the server's replies. The server can therefore provide sufficient information to the client to allow it to disambiguate between an erroneous or conflicting callback and a race condition.
For each client operation which might result in some sort of server callback, the server should "remember" the { slotid, sequenceid } pair of the client request until the slotid retirement rules allow the server to determine that the client has, in fact, seen the server's reply. Until the time the { slotid, sequenceid } request pair can be retired, any recalls of the associated object MUST carry an array of these referring identifiers (in the CB_SEQUENCE operation's arguments), for the benefit of the client. After this time, it is not necessary for the server to provide this information in related callbacks, since it is certain that a race condition can no longer occur.
The CB_SEQUENCE operation which begins each server callback carries a list of "referring" { slotid, sequenceid } tuples. If the client finds the request corresponding to the referring slotid and sequenced id be currently outstanding (i.e. the server's reply has not been seen by the client), it can determine that the callback has raced the reply, and act accordingly.
The client must not simply wait forever for the expected server reply to arrive on any of the session's operations channels, because it is possible that they will be delayed indefinitely. However, it should wait for a period of time, and if the time expires it can provide a more meaningful error such as NFS4ERR_DELAY.
[Comment.7] (We need to consider the clients' options here, and describe them... NFS4ERR_DELAY has been discussed as a legal reply to CB_RECALL?)
There are other scenarios under which callbacks may race replies, among them pnfs layout recalls, described in Section 12.3.5.3 (Serialization of Layout Operations) [Comment.8] (fill in the blanks w/others, etc...)
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Very large requests and replies may pose both buffer management issues (especially with RDMA) and reply cache issues. When the session is created, (Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid)) the client and server negotiate the maximum sized request they will send or process (ca_maxrequestsize), the maximum sized reply they will return or process (ca_maxresponsesize), and the maximum sized reply they will store in the reply cache (ca_maxresponsesize_cached).
If a request exceeds ca_maxrequestsize, the reply will have the status NFS4ERR_REQ_TOO_BIG. A replier may return NFS4ERR_REQ_TOO_BIG as the status for first operation (SEQUENCE or CB_SEQUENCE) in the request, or it may chose to return it on a subsequent operation.
If a reply exceeds ca_maxresponsesize, the reply will have the status NFS4ERR_REP_TOO_BIG. A replier may return NFS4ERR_REP_TOO_BIG as the status for first operation (SEQUENCE or CB_SEQUENCE) in the request, or it may chose to return it on a subsequent operation.
If sa_cachethis or csa_cachethis are TRUE, then the replier MUST cache a reply except if an error is returned by the SEQUENCE or CB_SEQUENCE operation (see Section 2.9.4.1.1 (Errors from SEQUENCE and CB_SEQUENCE)). If the reply exceeds ca_maxresponsesize_cached, (and sa_cachethis or csa_cachethis are TRUE) then the server MUST return NFS4ERR_REP_TOO_BIG_TO_CACHE. Even if NFS4ERR_REP_TOO_BIG_TO_CACHE (or any other error for that matter) is returned on a operation other than first operation (SEQUENCE or CB_SEQUENCE), then the reply MUST be cached if sa_cachethis or csa_cachethis are TRUE. For example, if a COMPOUND has eleven operations, including SEQUENCE, the fifth operation is a RENAME, and the tenth operation is a READ for one million bytes, server may return NFS4ERR_REP_TOO_BIG_TO_CACHE on the tenth operation. Since the server executed several operations, especially the non-idempotent RENAME, the client's request to cache the reply needs to be honored in order for correct operation of exactly once semantics. If the client retries the request, the server will have cached a reply that contains results for ten of the eleven requested operations, with the tenth operation having a status of NFS4ERR_REP_TOO_BIG_TO_CACHE.
A client needs to take care that when sending operations that change the current filehandle (except for PUTFH, PUTPUBFH, and PUTROOFFH) that it not exceed the maximum reply buffer before the GETFH operation. Otherwise the client will have to retry the operation that changed the current filehandle, in order obtain the desired filehandle. For the OPEN operation (see Section 16.16 (Operation 18: OPEN - Open a Regular File)), retry is not always available as an option. The following guidelines for the handling of filehandle changing operations are advised:
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Since the reply cache is bounded, it is practical for the server reply cache to persist across server reboots, and to be kept in stable storage (a client's reply cache for callbacks need not persist across client reboots unless the client intends for its session and other state to persist across reboots).
The CREATE_SESSION (see Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid) operation determines the persistence of the reply cache.
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A complete discussion of the operation of RPC-based protocols atop RDMA transports is in [RPCRDMA]. A discussion of the operation of NFSv4, including NFSv4.1 over RDMA is in [NFSDDP]. Where RDMA is considered, this specification assumes the use of such a layering; it addresses only the upper layer issues relevant to making best use of RPC/RDMA.
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RDMA requires its consumers to register memory and post buffers of a specific size and number for receive operations.
Registration of memory can be a relatively high-overhead operation, since it requires pinning of buffers, assignment of attributes (e.g. readable/writable), and initialization of hardware translation. Preregistration is desirable to reduce overhead. These registrations are specific to hardware interfaces and even to RDMA connection endpoints, therefore negotiation of their limits is desirable to manage resources effectively.
Following the basic registration, these buffers must be posted by the RPC layer to handle receives. These buffers remain in use by the RPC/NFSv4 implementation; the size and number of them must be known to the remote peer in order to avoid RDMA errors which would cause a fatal error on the RDMA connection.
NFSv4.1 manages slots as resources on a per session basis (see Section 2.9 (Session)), while RDMA connections manage credits on a per connection basis. This means that in order for a peer to send data over RDMA to a remote buffer, it has to have both an NFSv4.1 slot, and an RDMA credit.
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NFSv4.0 and all previous versions do not provide for any form of flow control; instead they rely on the windowing provided by transports like TCP to throttle requests. This does not work with RDMA, which provides no operation flow control and will terminate a connection in error when limits are exceeded. Limits such as maximum number of requests outstanding are therefore negotiated when a session is created (see the ca_maxrequests field in Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid)). These limits then provide the maxima each session's channels' connections must operate within. RDMA connections are managed within these limits as described in section 3.3 of [RPCRDMA]; if there are multiple RDMA connections, then the maximum requests for a channel will be divided among the RDMA connections. The limits may also be modified dynamically at the server's choosing by manipulating certain parameters present in each NFSv4.1 request. In addition, the CB_RECALL_SLOT callback operation (see Section 18.8 (Operation 10: CB_RECALL_SLOT - change flow control limits) can be issued by a server to a client to return RDMA credits to the server, thereby lowering the maximum number of requests a client can have outstanding to the server.
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Header padding is requested by each peer at session initiation (see the csa_headerpadsize argument to CREATE_SESSION in Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid)), and subsequently used by the RPC RDMA layer, as described in [RPCRDMA]. Zero padding is permitted.
Padding leverages the useful property that RDMA receives preserve alignment of data, even when they are placed into anonymous (untagged) buffers. If requested, client inline writes will insert appropriate pad bytes within the request header to align the data payload on the specified boundary. The client is encouraged to add sufficient padding (up to the negotiated size) so that the "data" field of the NFSv4.1 WRITE operation is aligned. Most servers can make good use of such padding, which allows them to chain receive buffers in such a way that any data carried by client requests will be placed into appropriate buffers at the server, ready for file system processing. The receiver's RPC layer encounters no overhead from skipping over pad bytes, and the RDMA layer's high performance makes the insertion and transmission of padding on the sender a significant optimization. In this way, the need for servers to perform RDMA Read to satisfy all but the largest client writes is obviated. An added benefit is the reduction of message round trips on the network - a potentially good trade, where latency is present.
The value to choose for padding is subject to a number of criteria. A primary source of variable-length data in the RPC header is the authentication information, the form of which is client-determined, possibly in response to server specification. The contents of COMPOUNDs, sizes of strings such as those passed to RENAME, etc. all go into the determination of a maximal NFSv4 request size and therefore minimal buffer size. The client must select its offered value carefully, so as not to overburden the server, and vice- versa. The payoff of an appropriate padding value is higher performance.
Sender gather:
|RPC Request|Pad bytes|Length| -> |User data...|
\------+---------------------/ \
\ \
\ Receiver scatter: \-----------+- ...
/-----+----------------\ \ \
|RPC Request|Pad|Length| -> |FS buffer|->|FS buffer|->...
In the above case, the server may recycle unused buffers to the next posted receive if unused by the actual received request, or may pass the now-complete buffers by reference for normal write processing. For a server which can make use of it, this removes any need for data copies of incoming data, without resorting to complicated end-to-end buffer advertisement and management. This includes most kernel-based and integrated server designs, among many others. The client may perform similar optimizations, if desired.
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Some RDMA transports (for example see [RDDP]), [Comment.9] (need xref) require a "streaming" (non-RDMA) phase, where ordinary traffic might flow before "stepping" up to RDMA mode, commencing RDMA traffic. Some RDMA transports start connections always in RDMA mode. NFSv4.1 allows, but does not assume, a streaming phase before RDMA mode. When a connection is bound to a session, the client and server negotiate whether the connection is used in RDMA or non-RDMA mode (see Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid) and Section 16.34 (Operation 41: BIND_CONN_TO_SESSION)).
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The session connection binding improves security over that provided by NFSv4.0 for the callback channel. The connection is client-initiated (see Section 16.34 (Operation 41: BIND_CONN_TO_SESSION)), and subject to the same firewall and routing checks as the operations channel. The connection cannot be hijacked by an attacker who connects to the client port prior to the intended server. At the client's option (see Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid) binding is fully authenticated before being activated (see Section 16.34 (Operation 41: BIND_CONN_TO_SESSION)). Traffic from the server over the callback channel is authenticated exactly as the client specifies (see Section 2.9.6.2 (Backchannel RPC Security)).
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When the NFSv4.1 client establishes the backchannel, it informs the server what security flavors and principals it must use when sending requests over the backchannel. If the security flavor is RPCSEC_GSS, the client expresses the principal in the form of an established RPCSEC_GSS context. The server is free to use any flavor/principal combination the server offers, but MUST NOT use unoffered combinations.
This way, the client does not have to provide a target GSS principal as it did with NFSv4.0, and the server does not have to implement an RPCSEC_GSS initiator as it did with NFSv4.0. [Comment.10] (xrefs)
The CREATE_SESSION (Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid)) and BACKCHANNEL_CTL (Section 16.33 (Operation 40: BACKCHANNEL_CTL - Backchannel control)) operations allow the client to specify flavor/principal combinations.
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Under some conditions, NFSv4.0 is vulnerable to a denial of service issue with respect to its state management.
The attack works via an unauthorized client faking an open_owner4, an open_owner/lock_owner pair, or stateid, combined with a seqid. The operation is sent to the NFSv4 server. The NFSv4 server accepts the state information, and as long as any status code from the result of this operation is not NFS4ERR_STALE_CLIENTID, NFS4ERR_STALE_STATEID, NFS4ERR_BAD_STATEID, NFS4ERR_BAD_SEQID, NFS4ERR_BADXDR, NFS4ERR_RESOURCE, or NFS4ERR_NOFILEHANDLE, the sequence number is incremented. When the authorized client issues an operation, it gets back NFS4ERR_BAD_SEQID, because its idea of the current sequence number is off by one. The authorized client's recovery options are pretty limited, with SETCLIENTID, followed by complete reclaim of state, which may or may not succeed completely. That qualifies as a denial of service attack.
If the client uses RPCSEC_GSS authentication and integrity, and every client maps each open_owner and lock_owner one and only one principal, and the server enforces this binding, then the conditions leading to vulnerability to the denial of service do not exist. One should keep in mind that if AUTH_SYS is being used, far simpler easier denial of service and other attacks are possible.
With NFSv4.1 sessions, the per-operation sequence number is ignored (see Section 8.13 (Vestigial Locking Infrastructure From V4.0)) therefore the NFSv4.0 denial of service vulnerability described above does not apply. However as described to this point in the specification, an attacker could forge the sessionid and issue a SEQUENCE with a slot id that he expects the legitimate client to use next. The legitimate client could then use the slotid with the same sequence number, and the server returns the attacker's result from the replay cache, thereby disrupting the legitimate client.
If we give each NFSv4.1 user their own session, and each user uses RPCSEC_GSS authentication and integrity, then the denial of service issue is solved, at the cost of additional per session state. The alternative NFSv4.1 specifies is described as follows.
Transport connections MUST be bound to a session by the client. The server MUST return an error to an operation (other than the operation that binds the connection to the session) that uses an unbound connection. As a simplification, the transport connection used by CREATE_SESSION (see Section 16.36 (Operation 43: CREATE_SESSION - Create New Session and Confirm Clientid)) is automatically bound to the session. Additional connections are bound to a session via BIND_CONN_TO_SESSION (see Section 16.34 (Operation 41: BIND_CONN_TO_SESSION)).
To prevent attackers from issuing BIND_CONN_TO_SESSION operations, the arguments to BIND_CONN_TO_SESSION include a digest of a shared secret called the secret session verifier (SSV) that only the client and server know. The digest is created via a one way, collision resistant hash function, making it intractable for the attacker to forge.
The SSV is sent to the server via SET_SSV (see Section 16.47 (Operation 54: SET_SSV)). To prevent eavesdropping, a SET_SSV for the SSV SHOULD be protected via RPCSEC_GSS with the privacy service. The SSV can be changed by the client at any time, by any principal. However several aspects of SSV changing prevent an attacker from engaging in a successful denial of service attack:
If a connection is disconnected, BIND_CONN_TO_SESSION is required to bind a connection to the session, even if the connection that was disconnected was the one CREATE_SESSION was created with.
If a client is assigned a machine principal then the client SHOULD use the machine principal's RPCSEC_GSS context to privacy protect the SSV from eavesdropping during the SET_SSV operation. If a machine principal is not being used, then the client MAY use the non-machine principal's RPCSEC_GSS context to privacy protect the SSV. The server MUST accept either type of principal. A client SHOULD change the SSV each time a new principal uses the session.
Here are the types of attacks that can be attempted by an attacker named Eve, and how the connection to session binding approach addresses each attack:
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The server has the primary obligation to monitor the state of backchannel resources that the client has created for the server (RPCSEC_GSS contexts and back channel connections). When these resources go away, the server takes action as specified in Section 2.9.8.2 (Events Requiring Server Action).
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The client has the following obligations in order to utilize the session:
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The client issues EXCHANGE_ID to establish a clientid.
The client uses the clientid to issue a CREATE_SESSION on a connection to the server. The results of CREATE_SESSION indicate whether the server will persist the session replay cache through a server reboot or not, and the client notes this for future reference.
The client SHOULD have specified connecting binding enforcement when the session was created. If so, the client SHOULD issue SET_SSV in the first COMPOUND after the session is created. If it is not using machine credentials, then each time a new principal goes to use the session, it SHOULD issue a SET_SSV again.
If the client wants to use delegations, layouts, directory notifications, or any other state that requires a callback channel, then it MUST add a connection to the backchannel if CREATE_SESSION did not already do so. The client creates a connection, and calls BIND_CONN_TO_SESSION to bind the connection to the session and the session's backchannel. If CREATE_SESSION did not already do so, the client MUST tell the server what security is required in order for the client to accept callbacks. The client does this via BACKCHANNEL_CTL.
If the client wants to use additional connections for the backchannel, then it MUST call BIND_CONN_TO_SESSION on each connection it wants to use with the session. If the client wants to use additional connections for the operation channel, then it MUST call BIND_CONN_TO_SESSION if it specified connection binding enforcement before using the connection.
At this point the client has reached a steady state as far as session use.
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The following events require client action to recover.
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If all RPCSEC_GSS contexts granted to by the client to the server for callback use have expired, the client MUST establish a new context via BACKCHANNEL_CTL. The sr_status field of SEQUENCE results indicates when callback contexts are nearly expired, or fully expired (see Section 16.46.4 (DESCRIPTION)).
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If the client loses the last connection of the session, then it MUST create a new connection, and if connecting binding enforcement was specified when the session was created, bind it to the session via BIND_CONN_TO_SESSION.
If there were requests outstanding at the time the of connection disconnect, then the client MUST retry the request, as described in Section 2.9.4.2 (Retry and Replay). Note that it is not necessary to retry requests over a connection with the same source network address or the same destination network address as the disconnected connection. As long as the sessionid, slotid, and sequenceid in the retry match that of the original request, the server will recognize the request as a retry if it did see the request prior to disconnect.
If the connection that was bound to the backchannel is lost, the client may need to reconnect, and use BIND_CONN_TO_SESSION, to give the connection to the backchannel. If the connection that was lost was the last one bound to the backchannel, the client MUST reconnect, and bind the connection to the session and backchannel. The server should indicate when it has no callback connection via the sr_status result from SEQUENCE.
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Via the sr_status result of the SEQUENCE operation or other means, the client will learn if some or all of the RPCSEC_GSS contexts it assigned to the backchannel have been lost. The client may need to use BACKCHANNEL_CTL to assign new contexts. It MUST assign new contexts if there are no more contexts.
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The server may lose a record of the session. Causes include:
Loss of replay cache is equivalent to loss of session. The server indicates loss of session to the client by returning NFS4ERR_BADSESSION on the next operation that uses the sessionid associated with the lost session.
After an event like a server reboot, the client may have lost its connections. The client assumes for the moment that the session has not been lost. It reconnects, and if it specified connecting binding enforcement when the session was created, it invokes BIND_CONN_TO_SESSION using the sessionid. Otherwise, it invokes SEQUENCE. If BIND_CONN_TO_SESSION or SEQUENCE returns NFS4ERR_BADSESSION, the client knows the session was lost. If the connection survives session loss, then the next SEQUENCE operation the client issues over the connection will get back NFS4ERR_BADSESSION. The client again knows the session was lost.
When the client detects session loss, it must call CREATE_SESSION to recover. Any non-idempotent operations that were in progress may have been performed on the server at the time of session loss. The client has no general way to recover from this.
Note that loss of session does not imply loss of lock, open, delegation, or layout state. Nor does loss of lock, open, delegation, or layout state imply loss of session state. [Comment.12] (Add reference to lock recovery section) . A session can survive a server reboot, but lock recovery may still be needed. The converse is also true.
It is possible CREATE_SESSION will fail with NFS4ERR_STALE_CLIENTID (for example the server reboots and does not preserve clientid state). If so, the client needs to call EXCHANGE_ID, followed by CREATE_SESSION.
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[Comment.13] (Dave Noveck requested this section; not sure what is needed here if this refers to failover to a replica. What are the session ramifications?)
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The following events require server action to recover.
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As described in Section 16.35 (Operation 42: EXCHANGE_ID - Instantiate Clientid), a rebooted client causes the server to delete any sessions it had.
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If a client crashes and never comes back, it will never issue EXCHANGE_ID with its old clientid. Thus the server has session state that will never be used again. After an extended period of time and if the server has resource constraints, it MAY destroy the old session.
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To the server, the extended network partition may be no different than a client crash with no reboot (see Section 2.9.8.2.2 (Client Crash with No Reboot)). Unless the server can discern that there is a network partition, it is free to treat the situation as if the client has crashed for good.
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If there were callback requests outstanding at the time the of a connection disconnect, then the server MUST retry the request, as described in Section 2.9.4.2 (Retry and Replay). Note that it is not necessary to retry requests over a connection with the same source network address or the same destination network address as the disconnected connection. As long as the sessionid, slotid, and sequenceid in the retry match that of the original request, the callback target will recognize the request as a retry if it did see the request prior to disconnect.
If the connection lost is the last one bound to the backchannel, then the server MUST indicate that in the sr_status field of the next SEQUENCE reply.
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The server SHOULD monitor when the last RPCSEC_GSS context assigned to the backchannel is near expiry (i.e. between one and two periods of lease time), and indicate so in the sr_status field of the next SEQUENCE reply. The server MUST indicate when the backchannel's last RPCSEC_GSS context has expired in the sr_status field of the next SEQUENCE reply.
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The syntax and semantics to describe the data types of the NFS version 4 protocol are defined in the XDR RFC4506 (Eisler, M., “XDR: External Data Representation Standard,” May 2006.) [3] and RPC RFC1831 (Srinivasan, R., “RPC: Remote Procedure Call Protocol Specification Version 2,” August 1995.) [4] documents. The next sections build upon the XDR data types to define types and structures specific to this protocol.
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These are the base NFSv4 data types.
| Data Type | Definition |
|---|---|
| int32_t | typedef int int32_t; |
| uint32_t | typedef unsigned int uint32_t; |
| int64_t | typedef hyper int64_t; |
| uint64_t | typedef unsigned hyper uint64_t; |
| attrlist4 | typedef opaque attrlist4<>; |
| Used for file/directory attributes | |
| bitmap4 | typedef uint32_t bitmap4<>; |
| Used in attribute array encoding. | |
| changeid4 | typedef uint64_t changeid4; |
| Used in definition of change_info | |
| clientid4 | typedef uint64_t clientid4; |
| Shorthand reference to client identification | |
| component4 | typedef utf8str_cs component4; |
| Represents path name components | |
| count4 | typedef uint32_t count4; |
| Various count parameters (READ, WRITE, COMMIT) | |
| length4 | typedef uint64_t length4; |
| Describes LOCK lengths | |
| linktext4 | typedef utf8str_cs linktext4; |
| Symbolic link contents | |
| mode4 | typedef uint32_t mode4; |
| Mode attribute data type | |
| nfs_cookie4 | typedef uint64_t nfs_cookie4; |
| Opaque cookie value for READDIR | |
| nfs_fh4 | typedef opaque nfs_fh4<NFS4_FHSIZE> |
| Filehandle definition; NFS4_FHSIZE is defined as 128 | |
| nfs_ftype4 | enum nfs_ftype4; |
| Various defined file types | |
| nfsstat4 | enum nfsstat4; |
| Return value for operations | |
| offset4 | typedef uint64_t offset4; |
| Various offset designations (READ, WRITE, LOCK, COMMIT) | |
| pathname4 | typedef component4 pathname4<>; |
| Represents path name for fs_locations | |
| qop4 | typedef uint32_t qop4; |
| Quality of protection designation in SECINFO | |
| sec_oid4 | typedef opaque sec_oid4<>; |
| Security Object Identifier The sec_oid4 data type is not really opaque. Instead contains an ASN.1 OBJECT IDENTIFIER as used by GSS-API in the mech_type argument to GSS_Init_sec_context. See RFC2743 (Linn, J., “Generic Security Service Application Program Interface Version 2, Update 1,” January 2000.) [8] for details. | |
| seqid4 | typedef uint32_t seqid4; |
| Sequence identifier used for file locking | |
| utf8string | typedef opaque utf8string<>; |
| UTF-8 encoding for strings | |
| utf8str_cis | typedef opaque utf8str_cis; |
| Case-insensitive UTF-8 string | |
| utf8str_cs | typedef opaque utf8str_cs; |
| Case-sensitive UTF-8 string | |
| utf8str_mixed | typedef opaque utf8str_mixed; |
| UTF-8 strings with a case sensitive prefix and a case insensitive suffix. | |
| verifier4 | typedef opaque verifier4[NFS4_VERIFIER_SIZE]; |
| Verifier used for various operations (COMMIT, CREATE, OPEN, READDIR, SETCLIENTID, SETCLIENTID_CONFIRM, WRITE) NFS4_VERIFIER_SIZE is defined as 8. |
End of Base Data Types
| Table 1 |
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struct nfstime4 {
int64_t seconds;
uint32_t nseconds;
}
The nfstime4 structure gives the number of seconds and nanoseconds since midnight or 0 hour January 1, 1970 Coordinated Universal Time (UTC). Values greater than zero for the seconds field denote dates after the 0 hour January 1, 1970. Values less than zero for the seconds field denote dates before the 0 hour January 1, 1970. In both cases, the nseconds field is to be added to the seconds field for the final time representation. For example, if the time to be represented is one-half second before 0 hour January 1, 1970, the seconds field would have a value of negative one (-1) and the nseconds fields would have a value of one-half second (500000000). Values greater than 999,999,999 for nseconds are considered invalid.
This data type is used to pass time and date information. A server converts to and from its local representation of time when processing time values, preserving as much accuracy as possible. If the precision of timestamps stored for a file system object is less than defined, loss of precision can occur. An adjunct time maintenance protocol is recommended to reduce client and server time skew.
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enum time_how4 {
SET_TO_SERVER_TIME4 = 0,
SET_TO_CLIENT_TIME4 = 1
};
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union settime4 switch (time_how4 set_it) {
case SET_TO_CLIENT_TIME4:
nfstime4 time;
default:
void;
};
The above definitions are used as the attribute definitions to set time values. If set_it is SET_TO_SERVER_TIME4, then the server uses its local representation of time for the time value.
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struct specdata4 {
uint32_t specdata1; /* major device number */
uint32_t specdata2; /* minor device number */
};
This data type represents additional information for the device file types NF4CHR and NF4BLK.
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struct fsid4 {
uint64_t major;
uint64_t minor;
};
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struct fs_location4 {
utf8str_cis server<>;
pathname4 rootpath;
};
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struct fs_locations4 {
pathname4 fs_root;
fs_location4 locations<>;
};
The fs_location4 and fs_locations4 data types are used for the fs_locations recommended attribute which is used for migration and replication support.
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struct fattr4 {
bitmap4 attrmask;
attrlist4 attr_vals;
};
The fattr4 structure is used to represent file and directory attributes.
The bitmap is a counted array of 32 bit integers used to contain bit values. The position of the integer in the array that contains bit n can be computed from the expression (n / 32) and its bit within that integer is (n mod 32).
0 1 +-----------+-----------+-----------+-- | count | 31 .. 0 | 63 .. 32 | +-----------+-----------+-----------+--
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struct change_info4 {
bool atomic;
changeid4 before;
changeid4 after;
};
This structure is used with the CREATE, LINK, REMOVE, RENAME operations to let the client know the value of the change attribute for the directory in which the target file system object resides.
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struct netaddr4 {
/* see struct rpcb in RFC1833 */
string r_netid<>; /* network id */
string r_addr<>; /* universal address */
};
The netaddr4 structure is used to identify TCP/IP based endpoints. The r_netid and r_addr fields are specified in RFC1833 (Srinivasan, R., “Binding Protocols for ONC RPC Version 2,” August 1995.) [22], but they are underspecified in RFC1833 (Srinivasan, R., “Binding Protocols for ONC RPC Version 2,” August 1995.) [22] as far as what they should look like for specific protocols.
For TCP over IPv4 and for UDP over IPv4, the format of r_addr is the US-ASCII string:
h1.h2.h3.h4.p1.p2
The prefix, "h1.h2.h3.h4", is the standard textual form for representing an IPv4 address, which is always four octets long. Assuming big-endian ordering, h1, h2, h3, and h4, are respectively, the first through fourth octets each converted to ASCII-decimal. Assuming big-endian ordering, p1 and p2 are, respectively, the first and second octets each converted to ASCII-decimal. For example, if a host, in big-endian order, has an address of 0x0A010307 and there is a service listening on, in big endian order, port 0x020F (decimal 527), then complete universal address is "10.1.3.7.2.15".
For TCP over IPv4 the value of r_netid is the string "tcp". For UDP over IPv4 the value of r_netid is the string "udp".
For TCP over IPv6 and for UDP over IPv6, the format of r_addr is the US-ASCII string:
x1:x2:x3:x4:x5:x6:x7:x8.p1.p2
The suffix "p1.p2" is the service port, and is computed the same way as with universal addresses for TCP and UDP over IPv4. The prefix, "x1:x2:x3:x4:x5:x6:x7:x8", is the standard textual form for representing an IPv6 address as defined in Section 2.2 of RFC1884 (Hinden, R. and S. Deering, “IP Version 6 Addressing Architecture,” December 1995.) [9]. Additionally, the two alternative forms specified in Section 2.2 of RFC1884 (Hinden, R. and S. Deering, “IP Version 6 Addressing Architecture,” December 1995.) [9] are also acceptable.
For TCP over IPv6 the value of r_netid is the string "tcp6". For UDP over IPv6 the value of r_netid is the string "udp6".
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typedef netaddr4 clientaddr4;
The clientaddr4 structure is used as part of the SETCLIENTID operation to either specify the address of the client that is using a clientid or as part of the callback registration.
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struct cb_client4 {
unsigned int cb_program;
netaddr4 cb_location;
};
This structure is used by the client to inform the server of its call back address; includes the program number and client address.
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struct nfs_client_id4 {
verifier4 verifier;
opaque id<NFS4_OPAQUE_LIMIT>
};
This structure is part of the arguments to the SETCLIENTID operation. NFS4_OPAQUE_LIMIT is defined as 1024.
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struct open_owner4 {
clientid4 clientid;
opaque owner<NFS4_OPAQUE_LIMIT>
};
This structure is used to identify the owner of open state. NFS4_OPAQUE_LIMIT is defined as 1024.
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struct lock_owner4 {
clientid4 clientid;
opaque owner<NFS4_OPAQUE_LIMIT>
};
This structure is used to identify the owner of file locking state. NFS4_OPAQUE_LIMIT is defined as 1024.
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struct open_to_lock_owner4 {
seqid4 open_seqid;
stateid4 open_stateid;
seqid4 lock_seqid;
lock_owner4 lock_owner;
};
This structure is used for the first LOCK operation done for an open_owner4. It provides both the open_stateid and lock_owner such that the transition is made from a valid open_stateid sequence to that of the new lock_stateid sequence. Using this mechanism avoids the confirmation of the lock_owner/lock_seqid pair since it is tied to established state in the form of the open_stateid/open_seqid.
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struct stateid4 {
uint32_t seqid;
opaque other[12];
};
This structure is used for the various state sharing mechanisms between the client and server. For the client, this data structure is read-only. The starting value of the seqid field is undefined. The server is required to increment the seqid field monotonically at each transition of the stateid. This is important since the client will inspect the seqid in OPEN stateids to determine the order of OPEN processing done by the server.
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enum layouttype4 {
LAYOUT_NFSV4_FILES = 1,
LAYOUT_OSD2_OBJECTS = 2,
LAYOUT_BLOCK_VOLUME = 3
};
A layout type specifies the layout being used. The implication is that clients have "layout drivers" that support one or more layout types. The file server advertises the layout types it supports through the LAYOUT_TYPES file system attribute. A client asks for layouts of a particular type in LAYOUTGET, and passes those layouts to its layout driver.
The layouttype4 structure is 32 bits in length. The range represented by the layout type is split into two parts. Types within the range 0x00000000-0x7FFFFFFF are globally unique and are assigned according to the description in Section 20.1 (Defining new layout types); they are maintained by IANA. Types within the range 0x80000000-0xFFFFFFFF are site specific and for "private use" only.
The LAYOUT_NFSV4_FILES enumeration specifies that the NFSv4 file layout type is to be used. The LAYOUT_OSD2_OBJECTS enumeration specifies that the object layout, as defined in [23] (Zelenka, J., Welch, B., and B. Halevy, “Object-based pNFS Operations,” July 2005.), is to be used. Similarly, the LAYOUT_BLOCK_VOLUME enumeration that the block/volume layout, as defined in [24] (Black, D., “pNFS Block/Volume Layout,” July 2005.), is to be used.
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typedef uint32_t deviceid4; /* 32-bit device ID */
Layout information includes device IDs that specify a storage device through a compact handle. Addressing and type information is obtained with the GETDEVICEINFO operation. A client must not assume that device IDs are valid across metadata server reboots. The device ID is qualified by the layout type and are unique per file system (FSID). This allows different layout drivers to generate device IDs without the need for co-ordination. See Section 12.3.1.4 (Device IDs) for more details.
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struct devlist_item4 {
deviceid4 dli_id;
opaque dli_device_addr<>;
};
An array of these values is returned by the GETDEVICELIST operation. They define the set of devices associated with a file system for the layout type specified in the GETDEVICELIST4args.
The device address is used to set up a communication channel with the storage device. Different layout types will require different types of structures to define how they communicate with storage devices. The opaque device_addr field must be interpreted based on the specified layout type.
This document defines the device address for the NFSv4 file layout (struct netaddr4 (netaddr4)), which identifies a storage device by network IP address and port number. This is sufficient for the clients to communicate with the NFSv4 storage devices, and may be sufficient for other layout types as well. Device types for object storage devices and block storage devices (e.g., SCSI volume labels) will be defined by their respective layout specifications.
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struct layout4 {
offset4 lo_offset;
length4 lo_length;
layoutiomode4 lo_iomode;
layouttype4 lo_type;
opaque lo_layout<>;
};
The layout4 structure defines a layout for a file. The layout type specific data is opaque within this structure and must be interepreted based on the layout type. Currently, only the NFSv4 file layout type is defined; see Section 12.4.2 (File Striping and Data Access) for its definition. Since layouts are sub-dividable, the offset and length together with the file's filehandle, the clientid, iomode, and layout type, identifies the layout.
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struct layoutupdate4 {
layouttype4 lou_type;
opaque lou_data<>;
};
The layoutupdate4 structure is used by the client to return 'updated' layout information to the metadata server at LAYOUTCOMMIT time. This structure provides a channel to pass layout type specific information back to the metadata server. E.g., for block/volume layout types this could include the list of reserved blocks that were written. The contents of the opaque lou_data argument are determined by the layout type and are defined in their context. The NFSv4 file-based layout does not use this structure, thus the update_data field should have a zero length.
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struct layouthint4 {
layouttype4 loh_type;
opaque loh_data<>;
};
The layouthint4 structure is used by the client to pass in a hint about the type of layout it would like created for a particular file. It is the structure specified by the FILE_LAYOUT_HINT attribute described below. The metadata server may ignore the hint, or may selectively ignore fields within the hint. This hint should be provided at create time as part of the initial attributes within OPEN. The NFSv4 file-based layout uses the "nfsv4_file_layouthint" structure as defined in Section 12.4.2 (File Striping and Data Access).
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enum layoutiomode4 {
LAYOUTIOMODE_READ = 1,
LAYOUTIOMODE_RW = 2,
LAYOUTIOMODE_ANY = 3
};
The iomode specifies whether the client intends to read or write (with the possibility of reading) the data represented by the layout. The ANY iomode MUST NOT be used for LAYOUTGET, however, it can be used for LAYOUTRETURN and LAYOUTRECALL. The ANY iomode specifies that layouts pertaining to both READ and RW iomodes are being returned or recalled, respectively. The metadata server's use of the iomode may depend on the layout type being used. The storage devices may validate I/O accesses against the iomode and reject invalid accesses.
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struct nfs_impl_id4 {
utf8str_cis nii_domain;
utf8str_cs nii_name;
nfstime4 nii_date;
};
This structure is used to identify client and server implementation detail. The nii_domain field is the DNS domain name that the implementer is associated with. The nii_name field is the product name of the implementation and is completely free form. It is encouraged that the nii_name be used to distinguish machine architecture, machine platforms, revisions, versions, and patch levels. The nii_date field is the timestamp of when the software instance was published or built.
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struct threshold_item4 {
layouttype4 thi_layout_type;
bitmap4 thi_hintset;
opaque thi_hintlist<>;
};
This structure contains a list of hints specific to a layout type for helping the client determine when it should issue I/O directly through the metadata server vs. the data servers. The hint structure consists of the layout type, a bitmap describing the set of hints supported by the server, they may differ based on the layout type, and a list of hints, whose structure is determined by the hintset bitmap. See the mdsthreshold attribute for more details.
The hintset is a bitmap of the following values:
| name | # | Data Type | Description |
|---|---|---|---|
| threshold4_read_size | 0 | length4 | The file size below which it is recommended to read data through the MDS. |
| threshold4_write_size | 1 | length4 | The file size below which it is recommended to write data through the MDS. |
| threshold4_read_iosize | 2 | length4 | For read I/O sizes below this threshold it is recommended to read data through the MDS |
| threshold4_write_iosize | 3 | length4 | For write I/O sizes below this threshold it is recommended to write data through the MDS |
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struct mdsthreshold4 {
threshold_item4 mth_hints<>;
};
This structure holds an array of threshold_item4 structures each of which is valid for a particular layout type. An array is necessary since a server can support multiple layout types for a single file.
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The filehandle in the NFS protocol is a per server unique identifier for a file system object. The contents of the filehandle are opaque to the client. Therefore, the server is responsible for translating the filehandle to an internal representation of the file system object.
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The operations of the NFS protocol are defined in terms of one or more filehandles. Therefore, the client needs a filehandle to initiate communication with the server. With the NFS version 2 protocol RFC1094 (Nowicki, B., “NFS: Network File System Protocol specification,” March 1989.) [17] and the NFS version 3 protocol RFC1813 (Callaghan, B., Pawlowski, B., and P. Staubach, “NFS Version 3 Protocol Specification,” June 1995.) [18], there exists an ancillary protocol to obtain this first filehandle. The MOUNT protocol, RPC program number 100005, provides the mechanism of translating a string based file system path name to a filehandle which can then be used by the NFS protocols.
The MOUNT protocol has deficiencies in the area of security and use via firewalls. This is one reason that the use of the public filehandle was introduced in RFC2054 (Callaghan, B., “WebNFS Client Specification,” October 1996.) [25] and RFC2055 (Callaghan, B., “WebNFS Server Specification,” October 1996.) [26]. With the use of the public filehandle in combination with the LOOKUP operation in the NFS version 2 and 3 protocols, it has been demonstrated that the MOUNT protocol is unnecessary for viable interaction between NFS client and server.
Therefore, the NFS version 4 protocol will not use an ancillary protocol for translation from string based path names to a filehandle. Two special filehandles will be used as starting points for the NFS client.
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The first of the special filehandles is the ROOT filehandle. The ROOT filehandle is the "conceptual" root of the file system name space at the NFS server. The client uses or starts with the ROOT filehandle by employing the PUTROOTFH operation. The PUTROOTFH operation instructs the server to set the "current" filehandle to the ROOT of the server's file tree. Once this PUTROOTFH operation is used, the client can then traverse the entirety of the server's file tree with the LOOKUP operation. A complete discussion of the server name space is in the section "NFS Server Name Space".
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The second special filehandle is the PUBLIC filehandle. Unlike the ROOT filehandle, the PUBLIC filehandle may be bound or represent an arbitrary file system object at the server. The server is responsible for this binding. It may be that the PUBLIC filehandle and the ROOT filehandle refer to the same file system object. However, it is up to the administrative software at the server and the policies of the server administrator to define the binding of the PUBLIC filehandle and server file system object. The client may not make any assumptions about this binding. The client uses the PUBLIC filehandle via the PUTPUBFH operation.
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In the NFS version 2 and 3 protocols, there was one type of filehandle with a single set of semantics. This type of filehandle is termed "persistent" in NFS Version 4. The semantics of a persistent filehandle remain the same as before. A new type of filehandle introduced in NFS Version 4 is the "volatile" filehandle, which attempts to accommodate certain server environments.
The volatile filehandle type was introduced to address server functionality or implementation issues which make correct implementation of a persistent filehandle infeasible. Some server environments do not provide a file system level invariant that can be used to construct a persistent filehandle. The underlying server file system may not provide the invariant or the server's file system programming interfaces may not provide access to the needed invariant. Volatile filehandles may ease the implementation of server functionality such as hierarchical storage management or file system reorganization or migration. However, the volatile filehandle increases the implementation burden for the client.
Since the client will need to handle persistent and volatile filehandles differently, a file attribute is defined which may be used by the client to determine the filehandle types being returned by the server.
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The filehandle contains all the information the server needs to distinguish an individual file. To the client, the filehandle is opaque. The client stores filehandles for use in a later request and can compare two filehandles from the same server for equality by doing a byte-by-byte comparison. However, the client MUST NOT otherwise interpret the contents of filehandles. If two filehandles from the same server are equal, they MUST refer to the same file. Servers SHOULD try to maintain a one-to-one correspondence between filehandles and files but this is not required. Clients MUST use filehandle comparisons only to improve performance, not for correct behavior. All clients need to be prepared for situations in which it cannot be determined whether two filehandles denote the same object and in such cases, avoid making invalid assumptions which might cause incorrect behavior. Further discussion of filehandle and attribute comparison in the context of data caching is presented in the section "Data Caching and File Identity".
As an example, in the case that two different path names when traversed at the server terminate at the same file system object, the server SHOULD return the same filehandle for each path. This can occur if a hard link is used to create two file names which refer to the same underlying file object and associated data. For example, if paths /a/b/c and /a/d/c refer to the same file, the server SHOULD return the same filehandle for both path names traversals.
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A persistent filehandle is defined as having a fixed value for the lifetime of the file system object to which it refers. Once the server creates the filehandle for a file system object, the server MUST accept the same filehandle for the object for the lifetime of the object. If the server restarts or reboots the NFS server must honor the same filehandle value as it did in the server's previous instantiation. Similarly, if the file system is migrated, the new NFS server must honor the same filehandle as the old NFS server.
The persistent filehandle will be become stale or invalid when the file system object is removed. When the server is presented with a persistent filehandle that refers to a deleted object, it MUST return an error of NFS4ERR_STALE. A filehandle may become stale when the file system containing the object is no longer available. The file system may become unavailable if it exists on removable media and the media is no longer available at the server or the file system in whole has been destroyed or the file system has simply been removed from the server's name space (i.e. unmounted in a UNIX environment).
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A volatile filehandle does not share the same longevity characteristics of a persistent filehandle. The server may determine that a volatile filehandle is no longer valid at many different points in time. If the server can definitively determine that a volatile filehandle refers to an object that has been removed, the server should return NFS4ERR_STALE to the client (as is the case for persistent filehandles). In all other cases where the server determines that a volatile filehandle can no longer be used, it should return an error of NFS4ERR_FHEXPIRED.
The mandatory attribute "fh_expire_type" is used by the client to determine what type of filehandle the server is providing for a particular file system. This attribute is a bitmask with the following values:
- FH4_PERSISTENT
- The value of FH4_PERSISTENT is used to indicate a persistent filehandle, which is valid until the object is removed from the file system. The server will not return NFS4ERR_FHEXPIRED for this filehandle. FH4_PERSISTENT is defined as a value in which none of the bits specified below are set.
- FH4_VOLATILE_ANY
- The filehandle may expire at any time, except as specifically excluded (i.e. FH4_NO_EXPIRE_WITH_OPEN).
- FH4_NOEXPIRE_WITH_OPEN
- May only be set when FH4_VOLATILE_ANY is set. If this bit is set, then the meaning of FH4_VOLATILE_ANY is qualified to exclude any expiration of the filehandle when it is open.
- FH4_VOL_MIGRATION
- The filehandle will expire as a result of a file system transition (migration or replication), in those case in which the continuity of filehandle use is not specified by handle class information within the fs_locations_info attribute. When this bit is set, clients without access to fs_locations_info information should assume filehandles will expire on file system transitions.
- FH4_VOL_RENAME
- The filehandle will expire during rename. This includes a rename by the requesting client or a rename by any other client. If FH4_VOL_ANY is set, FH4_VOL_RENAME is redundant.
Servers which provide volatile filehandles that may expire while open (i.e. if FH4_VOL_MIGRATION or FH4_VOL_RENAME is set or if FH4_VOLATILE_ANY is set and FH4_NOEXPIRE_WITH_OPEN not set), should deny a RENAME or REMOVE that would affect an OPEN file of any of the components leading to the OPEN file. In addition, the server should deny all RENAME or REMOVE requests during the grace period upon server restart.
Servers which provide volatile filehandles that may expire while open require special care as regards handling of RENAMESs and REMOVEs. This situation can arise if FH4_VOL_MIGRATION or FH4_VOL_RENAME is set, if FH4_VOLATILE_ANY is set and FH4_NOEXPIRE_WITH_OPEN not set, or if a non-readonly file system has a transition target in a different handle class. In these cases, the server should deny a RENAME or REMOVE that would affect an OPEN file of any of the components leading to the OPEN file. In addition, the server should deny all RENAME or REMOVE requests during the grace period, in order to make sure that reclaims of files where filehandles may have expired do not do a reclaim for the wrong file.
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A volatile filehandle, while opaque to the client could contain:
[volatile bit = 1 | server boot time | slot | generation number]
When the client presents a volatile filehandle, the server makes the following checks, which assume that the check for the volatile bit has passed. If the server boot time is less than the current server boot time, return NFS4ERR_FHEXPIRED. If slot is out of range, return NFS4ERR_BADHANDLE. If the generation number does not match, return NFS4ERR_FHEXPIRED.
When the server reboots, the table is gone (it is volatile).
If volatile bit is 0, then it is a persistent filehandle with a different structure following it.
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If possible, the client SHOULD recover from the receipt of an NFS4ERR_FHEXPIRED error. The client must take on additional responsibility so that it may prepare itself to recover from the expiration of a volatile filehandle. If the server returns persistent filehandles, the client does not need these additional steps.
For volatile filehandles, most commonly the client will need to store the component names leading up to and including the file system object in question. With these names, the client should be able to recover by finding a filehandle in the name space that is still available or by starting at the root of the server's file system name space.
If the expired filehandle refers to an object that has been removed from the file system, obviously the client will not be able to recover from the expired filehandle.
It is also possible that the expired filehandle refers to a file that has been renamed. If the file was renamed by another client, again it is possible that the original client will not be able to recover. However, in the case that the client itself is renaming the file and the file is open, it is possible that the client may be able to recover. The client can determine the new path name based on the processing of the rename request. The client can then regenerate the new filehandle based on the new path name. The client could also use the compound operation mechanism to construct a set of operations like:
RENAME A B
LOOKUP B
GETFH
Note that the COMPOUND procedure does not provide atomicity. This example only reduces the overhead of recovering from an expired filehandle.
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To meet the requirements of extensibility and increased interoperability with non-UNIX platforms, attributes must be handled in a flexible manner. The NFS version 3 fattr3 structure contains a fixed list of attributes that not all clients and servers are able to support or care about. The fattr3 structure can not be extended as new needs arise and it provides no way to indicate non-support. With the NFS version 4 protocol, the client is able query what attributes the server supports and construct requests with only those supported attributes (or a subset thereof).
To this end, attributes are divided into three groups: mandatory, recommended, and named. Both mandatory and recommended attributes are supported in the NFS version 4 protocol by a specific and well-defined encoding and are identified by number. They are requested by setting a bit in the bit vector sent in the GETATTR request; the server response includes a bit vector to list what attributes were returned in the response. New mandatory or recommended attributes may be added to the NFS protocol between major revisions by publishing a standards-track RFC which allocates a new attribute number value and defines the encoding for the attribute. See the section "Minor Versioning" for further discussion.
Named attributes are accessed by the new OPENATTR operation, which accesses a hidden directory of attributes associated with a file system object. OPENATTR takes a filehandle for the object and returns the filehandle for the attribute hierarchy. The filehandle for the named attributes is a directory object accessible by LOOKUP or READDIR and contains files whose names represent the named attributes and whose data bytes are the value of the attribute. For example:
| LOOKUP | "foo" | ; look up file |
| GETATTR | attrbits | |
| OPENATTR | ; access foo's named attributes | |
| LOOKUP | "x11icon" | ; look up specific attribute |
| READ | 0,4096 | ; read stream of bytes |
Named attributes are intended for data needed by applications rather than by an NFS client implementation. NFS implementors are strongly encouraged to define their new attributes as recommended attributes by bringing them to the IETF standards-track process.
The set of attributes which are classified as mandatory is deliberately small since servers must do whatever it takes to support them. A server should support as many of the recommended attributes as possible but by their definition, the server is not required to support all of them. Attributes are deemed mandatory if the data is both needed by a large number of clients and is not otherwise reasonably computable by the client when support is not provided on the server.
Note that the hidden directory returned by OPENATTR is a convenience for protocol processing. The client should not make any assumptions about the server's implementation of named attributes and whether the underlying file system at the server has a named attribute directory or not. Therefore, operations such as SETATTR and GETATTR on the named attribute directory are undefined.
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These MUST be supported by every NFS version 4 client and server in order to ensure a minimum level of interoperability. The server must store and return these attributes and the client must be able to function with an attribute set limited to these attributes. With just the mandatory attributes some client functionality may be impaired or limited in some ways. A client may ask for any of these attributes to be returned by setting a bit in the GETATTR request and the server must return their value.
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These attributes are understood well enough to warrant support in the NFS version 4 protocol. However, they may not be supported on all clients and servers. A client may ask for any of these attributes to be returned by setting a bit in the GETATTR request but must handle the case where the server does not return them. A client may ask for the set of attributes the server supports and should not request attributes the server does not support. A server should be tolerant of requests for unsupported attributes and simply not return them rather than considering the request an error. It is expected that servers will support all attributes they comfortably can and only fail to support attributes which are difficult to support in their operating environments. A server should provide attributes whenever they don't have to "tell lies" to the client. For example, a file modification time should be either an accurate time or should not be supported by the server. This will not always be comfortable to clients but the client is better positioned decide whether and how to fabricate or construct an attribute or whether to do without the attribute.
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These attributes are not supported by direct encoding in the NFS Version 4 protocol but are accessed by string names rather than numbers and correspond to an uninterpreted stream of bytes which are stored with the file system object. The name space for these attributes may be accessed by using the OPENATTR operation. The OPENATTR operation returns a filehandle for a virtual "attribute directory" and further perusal of the name space may be done using READDIR and LOOKUP operations on this filehandle. Named attributes may then be examined or changed by normal READ and WRITE and CREATE operations on the filehandles returned from READDIR and LOOKUP. Named attributes may have attributes.
It is recommended that servers support arbitrary named attributes. A client should not depend on the ability to store any named attributes in the server's file system. If a server does support named attributes, a client which is also able to handle them should be able to copy a file's data and meta-data with complete transparency from one location to another; this would imply that names allowed for regular directory entries are valid for named attribute names as well.
Names of attributes will not be controlled by this document or other IETF standards track documents. See the section "IANA Considerations" for further discussion.
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Each of the Mandatory and Recommended attributes can be classified in one of three categories: per server, per file system, or per file system object. Note that it is possible that some per file system attributes may vary within the file system. See the "homogeneous" attribute for its definition. Note that the attributes time_access_set and time_modify_set are not listed in this section because they are write-only attributes corresponding to time_access and time_modify, and are used in a special instance of SETATTR.
lease_time, send_impl_id, recv_impl_id
supp_attr, fh_expire_type, link_support, symlink_support, unique_handles, aclsupport, cansettime, case_insensitive, case_preserving, chown_restricted, files_avail, files_free, files_total, fs_locations, homogeneous, maxfilesize, maxname, maxread, maxwrite, no_trunc, space_avail, space_free, space_total, time_delta, fs_layout_type
type, change, size, named_attr, fsid, rdattr_error, filehandle, ACL, archive, fileid, hidden, maxlink, mimetype, mode, numlinks, owner, owner_group, rawdev, space_used, system, time_access, time_backup, time_create, time_metadata, time_modify, mounted_on_fileid, layout_type, layout_hint, layout_blksize, layout_alignment
For quota_avail_hard, quota_avail_soft, and quota_used see their definitions below for the appropriate classification.
| TOC |
| name | # | Data Type | Access | Description |
|---|---|---|---|---|
| supp_attr | 0 | bitmap | READ | The bit vector which would retrieve all mandatory and recommended attributes that are supported for this object. The scope of this attribute applies to all objects with a matching fsid. |
| type | 1 | nfs4_ftype | READ | The type of the object (file, directory, symlink, etc.) |
| fh_expire_type | 2 | uint32 | READ | Server uses this to specify filehandle expiration behavior to the client. See the section "Filehandles" for additional description. |
| change | 3 | uint64 | READ | A value created by the server that the client can use to determine if file data, directory contents or attributes of the object have been modified. The server may return the object's time_metadata attribute for this attribute's value but only if the file system object can not be updated more frequently than the resolution of time_metadata. |
| size | 4 | uint64 | R/W | The size of the object in bytes. |
| link_support | 5 | bool | READ | True, if the object's file system supports hard links. |
| symlink_support | 6 | bool | READ | True, if the object's file system supports symbolic links. |
| named_attr | 7 | bool | READ | True, if this object has named attributes. In other words, object has a non-empty named attribute directory. |
| fsid | 8 | fsid4 | READ | Unique file system identifier for the file system holding this object. fsid contains major and minor components each of which are uint64. |
| unique_handles | 9 | bool | READ | True, if two distinct filehandles guaranteed to refer to two different file system objects. |
| lease_time | 10 | nfs_lease4 | READ | Duration of leases at server in seconds. |
| rdattr_error | 11 | enum | READ | Error returned from getattr during readdir. |
| filehandle | 19 | nfs_fh4 | READ | The filehandle of this object (primarily for readdir requests). |
| TOC |
| name | # | Data Type | Access | Description |
|---|---|---|---|---|
| ACL | 12 | nfsace4<> | R/W | The access control list for the object. |
| aclsupport | 13 | uint32 | READ | Indicates what types of ACLs are supported on the current file system. |
| archive | 14 | bool | R/W | True, if this file has been archived since the time of last modification (deprecated in favor of time_backup). |
| cansettime | 15 | bool | READ | True, if the server able to change the times for a file system object as specified in a SETATTR operation. |
| case_insensitive | 16 | bool | READ | True, if filename comparisons on this file system are case insensitive. |
| case_preserving | 17 | bool | READ | True, if filename case on this file system are preserved. |
| chown_restricted | 18 | bool | READ | If TRUE, the server will reject any request to change either the owner or the group associated with a file if the caller is not a privileged user (for example, "root" in UNIX operating environments or in Windows 2000 the "Take Ownership" privilege). |
| dir_notif_delay | 56 | nfstime4 | READ | notification delays on directory attributes |
| dirent_ notif_delay | 57 | nfstime4 | READ | notification delays on child attributes |
| fileid | 20 | uint64 | READ | A number uniquely identifying the file within the file system. |
| files_avail | 21 | uint64 | READ | File slots available to this user on the file system containing this object - this should be the smallest relevant limit. |
| files_free | 22 | uint64 | READ | Free file slots on the file system containing this object - this should be the smallest relevant limit. |
| files_total | 23 | uint64 | READ | Total file slots on the file system containing this object. |
| fs_absent | 60 | bool | READ | Is current file system present or absent. |
| fs_layout_type | 62 | layouttype4<> | READ | Layout types available for the file system. |
| fs_locations | 24 | fs_locations | READ | Locations where this file system may be found. If the server returns NFS4ERR_MOVED as an error, this attribute MUST be supported. |
| fs_locations_info | 67 | READ | Full function file system location. | |
| fs_status | 61 | fs4_status | READ | Generic file system type information. |
| hidden | 25 | bool | R/W | True, if the file is considered hidden with respect to the Windows API? |
| homogeneous | 26 | bool | READ | True, if this object's file system is homogeneous, i.e. are per file system attributes the same for all file system's objects. |
| layout_alignment | 66 | uint32_t | READ | Preferred alignment for layout related I/O. |
| layout_blksize | 65 | uint32_t | READ | Preferred block size for layout related I/O. |
| layout_hint | 63 | layouthint4 | WRITE | Client specified hint for file layout. |
| layout_type | 64 | layouttype4<> | READ | Layout types available for the file. |
| maxfilesize | 27 | uint64 | READ | Maximum supported file size for the file system of this object. |
| maxlink | 28 | uint32 | READ | Maximum number of links for this object. |
| maxname | 29 | uint32 | READ | Maximum filename size supported for this object. |
| maxread | 30 | uint64 | READ | Maximum read size supported for this object. |
| maxwrite | 31 | uint64 | READ | Maximum write size supported for this object. This attribute SHOULD be supported if the file is writable. Lack of this attribute can lead to the client either wasting bandwidth or not receiving the best performance. |
| mdsthreshold | 68 | mdsthreshold4 | READ | Hint to client as to when to write through the pnfs metadata server. |
| mimetype | 32 | utf8<> | R/W | MIME body type/subtype of this object. |
| mode | 33 | mode4 | R/W | UNIX-style mode and permission bits for this object. |
| mounted_on_fileid | 55 | uint64 | READ | Like fileid, but if the target filehandle is the root of a file system return the fileid of the underlying directory. |
| no_trunc | 34 | bool | READ | True, if a name longer than name_max is used, an error be returned and name is not truncated. |
| numlinks | 35 | uint32 | READ | Number of hard links to this object. |
| owner | 36 | utf8<> | R/W | The string name of the owner of this object. |
| owner_group | 37 | utf8<> | R/W | The string name of the group ownership of this object. |
| quota_avail_hard | 38 | uint64 | READ | For definition see "Quota Attributes" section below. |
| quota_avail_soft | 39 | uint64 | READ | For definition see "Quota Attributes" section below. |
| quota_used | 40 | uint64 | READ | For definition see "Quota Attributes" section below. |
| rawdev | 41 | specdata4 | READ | Raw device identifier. UNIX device major/minor node information. If the value of type is not NF4BLK or NF4CHR, the value return SHOULD NOT be considered useful. |
| recv_impl_id | 59 | impl_ident4 | READ | Client obtains the server's implementation identity via GETATTR. |
| retentevt_get | 71 | retention_get4 | READ | Get the event-based retention duration, and if enabled, the event-based retention begin time of the file object. GETATTR use only. |
| retentevt_set | 72 | retention_set4 | WRITE | Set the event-based retention duration, and optionally enable event-based retention on the file object. SETATTR use only. |
| retention_get | 69 | retention_get4 | READ | Get the retention duration, and if enabled, the retention begin time of the file object. GETATTR use only. |
| retention_hold | 69 | uint64_t | R/W | Get or set administrative retention holds, one hold per bit position. |
| retention_set | 70 | retention_set4 | WRITE | Set the retention duration, and optionally enable retention on the file object. SETATTR use only. |
| send_impl_id | 58 | impl_ident4 | WRITE | Client provides server with its implementation identity via SETATTR. |
| space_avail | 42 | uint64 | READ | Disk space in bytes available to this user on the file system containing this object - this should be the smallest relevant limit. |
| space_free | 43 | uint64 | READ | Free disk space in bytes on the file system containing this object - this should be the smallest relevant limit. |
| space_total | 44 | uint64 | READ | Total disk space in bytes on the file system containing this object. |
| space_used | 45 | uint64 | READ | Number of file system bytes allocated to this object. |
| system | 46 | bool | R/W | True, if this file is a "system" file with respect to the Windows API? |
| time_access | 47 | nfstime4 | READ | The time of last access to the object by a read that was satisfied by the server. |
| time_access_set | 48 | settime4 | WRITE | Set the time of last access to the object. SETATTR use only. |
| time_backup | 49 | nfstime4 | R/W | The time of last backup of the object. |
| time_create | 50 | nfstime4 | R/W | The time of creation of the object. This attribute does not have any relation to the traditional UNIX file attribute "ctime" or "change time". |
| time_delta | 51 | nfstime4 | READ | Smallest useful server time granularity. |
| time_metadata | 52 | nfstime4 | READ | The time of last meta-data modification of the object. |
| time_modify | 53 | nfstime4 | READ | The time of last modification to the object. |
| time_modify_set | 54 | settime4 | WRITE | Set the time of last modification to the object. SETATTR use only. |
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As defined above, the time_access attribute represents the time of last access to the object by a read that was satisfied by the server. The notion of what is an "access" depends on server's operating environment and/or the server's file system semantics. For example, for servers obeying POSIX semantics, time_access would be updated only by the READLINK, READ, and READDIR operations and not any of the operations that modify the content of the object. Of course, setting the corresponding time_access_set attribute is another way to modify the time_access attribute.
Whenever the file object resides on a writable file system, the server should make best efforts to record time_access into stable storage. However, to mitigate the performance effects of doing so, and most especially whenever the server is satisfying the read of the object's content from its cache, the server MAY cache access time updates and lazily write them to stable storage. It is also acceptable to give administrators of the server the option to disable time_access updates.
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The recommended attributes "owner" and "owner_group" (and also users and groups within the "acl" attribute) are represented in terms of a UTF-8 string. To avoid a representation that is tied to a particular underlying implementation at the client or server, the use of the UTF-8 string has been chosen. Note that section 6.1 of RFC2624 (Shepler, S., “NFS Version 4 Design Considerations,” June 1999.) [27] provides additional rationale. It is expected that the client and server will have their own local representation of owner and owner_group that is used for local storage or presentation to the end user. Therefore, it is expected that when these attributes are transferred between the client and server that the local representation is translated to a syntax of the form "user@dns_domain". This will allow for a client and server that do not use the same local representation the ability to translate to a common syntax that can be interpreted by both.
Similarly, security principals may be represented in different ways by different security mechanisms. Servers normally translate these representations into a common format, generally that used by local storage, to serve as a means of identifying the users corresponding to these security principals. When these local identifiers are translated to the form of the owner attribute, associated with files created by such principals they identify, in a common format, the users associated with each corresponding set of security principals.
The translation used to interpret owner and group strings is not specified as part of the protocol. This allows various solutions to be employed. For example, a local translation table may be consulted that maps between a numeric id to the user@dns_domain syntax. A name service may also be used to accomplish the translation. A server may provide a more general service, not limited by any particular translation (which would only translate a limited set of possible strings) by storing the owner and owner_group attributes in local storage without any translation or it may augment a translation method by storing the entire string for attributes for which no translation is available while using the local representation for those cases in which a translation is available.
Servers that do not provide support for all possible values of the owner and owner_group attributes, should return an error (NFS4ERR_BADOWNER) when a string is presented that has no translation, as the value to be set for a SETATTR of the owner, owner_group, or acl attributes. When a server does accept an owner or owner_group value as valid on a SETATTR (and similarly for the owner and group strings in an acl), it is promising to return that same string when a corresponding GETATTR is done. Configuration changes and ill-constructed name translations (those that contain aliasing) may make that promise impossible to honor. Servers should make appropriate efforts to avoid a situation in which these attributes have their values changed when no real change to ownership has occurred.
The "dns_domain" portion of the owner string is meant to be a DNS domain name. For example, user@ietf.org. Servers should accept as valid a set of users for at least one domain. A server may treat other domains as having no valid translations. A more general service is provided when a server is capable of accepting users for multiple domains, or for all domains, subject to security constraints.
In the case where there is no translation available to the client or server, the attribute value must be constructed without the "@". Therefore, the absence of the @ from the owner or owner_group attribute signifies that no translation was available at the sender and that the receiver of the attribute should not use that string as a basis for translation into its own internal format. Even though the attribute value can not be translated, it may still be useful. In the case of a client, the attribute string may be used for local display of ownership.
To provide a greater degree of compatibility with previous versions of NFS (i.e. v2 and v3), which identified users and groups by 32-bit unsigned uid's and gid's, owner and group strings that consist of decimal numeric values with no leading zeros can be given a special interpretation by clients and servers which choose to provide such support. The receiver may treat such a user or group string as representing the same user as would be represented by a v2/v3 uid or gid having the corresponding numeric value. A server is not obligated to accept such a string, but may return an NFS4ERR_BADOWNER instead. To avoid this mechanism being used to subvert user and group translation, so that a client might pass all of the owners and groups in numeric form, a server SHOULD return an NFS4ERR_BADOWNER error when there is a valid translation for the user or owner designated in this way. In that case, the client must use the appropriate name@domain string and not the special form for compatibility.
The owner string "nobody" may be used to designate an anonymous user, which will be associated with a file created by a security principal that cannot be mapped through normal means to the owner attribute.
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With respect to the case_insensitive and case_preserving attributes, each UCS-4 character (which UTF-8 encodes) has a "long descriptive name" RFC1345 (Simonsen, K., “Character Mnemonics and Character Sets,” June 1992.) [28] which may or may not included the word "CAPITAL" or "SMALL". The presence of SMALL or CAPITAL allows an NFS server to implement unambiguous and efficient table driven mappings for case insensitive comparisons, and non-case-preserving storage. For general character handling and internationalization issues, see the section "Internationalization".
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For the attributes related to file system quotas, the following definitions apply:
- quota_avail_soft
- The value in bytes which represents the amount of additional disk space that can be allocated to this file or directory before the user may reasonably be warned. It is understood that this space may be consumed by allocations to other files or directories though there is a rule as to which other files or directories.
- quota_avail_hard
- The value in bytes which represent the amount of additional disk space beyond the current allocation that can be allocated to this file or directory before further allocations will be refused. It is understood that this space may be consumed by allocations to other files or directories.
- quota_used
- The value in bytes which represent the amount of disc space used by this file or directory and possibly a number of other similar files or directories, where the set of "similar" meets at least the criterion that allocating space to any file or directory in the set will reduce the "quota_avail_hard" of every other file or directory in the set.
Note that there may be a number of distinct but overlapping sets of files or directories for which a quota_used value is maintained. E.g. "all files with a given owner", "all files with a given group owner". etc.
The server is at liberty to choose any of those sets but should do so in a repeatable way. The rule may be configured per file system or may be "choose the set with the smallest quota".
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UNIX-based operating environments connect a file system into the namespace by connecting (mounting) the file system onto the existing file object (the mount point, usually a directory) of an existing file system. When the mount point's parent directory is read via an API like readdir(), the return results are directory entries, each with a component name and a fileid. The fileid of the mount point's directory entry will be different from the fileid that the stat() system call returns. The stat() system call is returning the fileid of the root of the mounted file system, whereas readdir() is returning the fileid stat() would have returned before any file systems were mounted on the mount point.
Unlike NFS version 3, NFS version 4 allows a client's LOOKUP request to cross other file systems. The client detects the file system crossing whenever the filehandle argument of LOOKUP has an fsid attribute different from that of the filehandle returned by LOOKUP. A UNIX-based client will consider this a "mount point crossing". UNIX has a legacy scheme for allowing a process to determine its current working directory. This relies on readdir() of a mount point's parent and stat() of the mount point returning fileids as previously described. The mounted_on_fileid attribute corresponds to the fileid that readdir() would have returned as described previously.
While the NFS version 4 client could simply fabricate a fileid corresponding to what mounted_on_fileid provides (and if the server does not support mounted_on_fileid, the client has no choice), there is a risk that the client will generate a fileid that conflicts with one that is already assigned to another object in the file system. Instead, if the server can provide the mounted_on_fileid, the potential for client operational problems in this area is eliminated.
If the server detects that there is no mounted point at the target file object, then the value for mounted_on_fileid that it returns is the same as that of the fileid attribute.
The mounted_on_fileid attribute is RECOMMENDED, so the server SHOULD provide it if possible, and for a UNIX-based server, this is straightforward. Usually, mounted_on_fileid will be requested during a READDIR operation, in which case it is trivial (at least for UNIX-based servers) to return mounted_on_fileid since it is equal to the fileid of a directory entry returned by readdir(). If mounted_on_fileid is requested in a GETATTR operation, the server should obey an invariant that has it returning a value that is equal to the file object's entry in the object's parent directory, i.e. what readdir() would have returned. Some operating environments allow a series of two or more file systems to be mounted onto a single mount point. In this case, for the server to obey the aforementioned invariant, it will need to find the base mount point, and not the intermediate mount points.
| TOC |
These recommended attributes are used to identify the client and server. In the case of the send_impl_id attribute, the client sends its nfs_impl_id4. In the case of the recv_impl_id attribute, the client receives the server's nfs_impl_id4 value.
Access to this identification information can be most useful at both client and server. Being able to identify specific implementations can help in planning by administrators or implementors. For example, diagnostic software may extract this information in an attempt to identify interoperability problems, performance workload behaviors or general usage statistics. Since the intent of having access to this information is for planning or general diagnosis only, the client and server MUST NOT interpret this implementation identity information in a way that affects interoperational behavior of the implementation. The reason is the if clients and servers did such a thing, they might use fewer capabilities of the protocol than the peer can support, or the client and server might refuse to interoperate.
Because it is likely some implementations will violate the protocol specification and interpret the identity information, implementations MUST allow the users of the NFSv4 client and server to set the contents of the sent nfs_impl_id structure to any value.
Even though these attributes are RECOMMENDED, if the server supports one of them it MUST support the other.
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This attribute applies to a file system and indicates what layout types are supported by the file system. We expect this attribute to be queried when a client encounters a new fsid. This attribute is used by the client to determine if it has applicable layout drivers.
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This attribute indicates the particular layout type(s) used for a file. This is for informational purposes only. The client needs to use the LAYOUTGET operation in order to get enough information (e.g., specific device information) in order to perform I/O.
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This attribute may be set on newly created files to influence the metadata server's choice for the file's layout. It is suggested that this attribute is set as one of the initial attributes within the OPEN call. The metadata server may ignore this attribute. This attribute is a sub-set of the layout structure returned by LAYOUTGET. For example, instead of specifying particular devices, this would be used to suggest the stripe width of a file. It is up to the server implementation to determine which fields within the layout it uses.
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This attribute acts as a hint to the client to help it determine when it is more efficient to issue read and write requests to the metadata server vs. the data server. Two types of thresholds are described: file size thresholds and I/O size thresholds. If a file's size is smaller than the file size threshold, data accesses should be issued to the metadata server. If an I/O is below the I/O size threshold, the I/O should be issued to the metadata server. Each threshold can be specified independently for read and write requests. For either threshold type, a value of 0 indicates no read or write should be issued to the metadata server, while a value of all 1s indicates all reads or writes should be issued to the metadata server.
The attribute is available on a per filehandle basis. If the current filehandle refers to a non-pNFS file or directory, the metadata server should return an attribute that is representative of the filehandle's file system. It is suggested that this attribute is queried as part of the OPEN operation. Due to dynamic system changes, the client should not assume that the attribute will remain constant for any specific time period, thus it should be periodically refreshed.
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Retention is a concept whereby a file object can be placed in an immutable, undeletable, unrenamable state for a fixed or infinite duration of time. Once in this "retained" state, the file cannot be moved out of the state until the duration of retention has been reached.
When retention is enabled, retention MUST extend to the data of the file, and the name of file. The server MAY extend retention any other property of the file, including any subset of mandatory, recommended, and named attributes, with the exceptions noted in this section.
Servers MAY support or not support retention on any file object type.
There are five retention attributes:
const RET4_DURATION_INFINITE = 0xffffffffffffffff;
struct retention_get4 {
uint64_t rg_duration;
nfstime4 rg_begin_time<1>;
};
struct retention_set4 {
bool rs_enable;
uint64_t rs_duration<1>;
};
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Access Control Lists (ACLs) are a file attribute that specify fine grained access control. This chapter covers the "acl", "aclsupport", and "mode" file attributes, and their interactions.
| TOC |
ACLs and modes represent two well established but different models for specifying permissions. This chapter specifies requirements that attempt to meet the following goals:
| TOC |
| TOC |
The NFS version 4 ACL attribute is an array of access control entries (ACEs). Although the client can read and write the ACL attribute, the server is responsible for using the ACL to perform access control. The client can use the OPEN or ACCESS operations to check access without modifying or reading data or metadata.
The NFS ACE attribute is defined as follows:
typedef uint32_t acetype4;
typedef uint32_t aceflag4;
typedef uint32_t acemask4;
struct nfsace4 {
acetype4 type;
aceflag4 flag;
acemask4 access_mask;
utf8str_mixed who;
};
To determine if a request succeeds, the server processes each nfsace4 entry in order. Only ACEs which have a "who" that matches the requester are considered. Each ACE is processed until all of the bits of the requester's access have been ALLOWED. Once a bit (see below) has been ALLOWED by an ACCESS_ALLOWED_ACE, it is no longer considered in the processing of later ACEs. If an ACCESS_DENIED_ACE is encountered where the requester's access still has unALLOWED bits in common with the "access_mask" of the ACE, the request is denied. When the ACL is fully processed, if there are bits in the requester's mask that have not been ALLOWED or DENIED, access is denied.
Unlike the ALLOW and DENY ACE types, the ALARM and AUDIT ACE types do not affect a requester's access, and instead are for triggering events as a result of a requester's access attempt. Therefore, all AUDIT and ALARM ACEs are processed until end of the ACL.
The NFS version 4 ACL model is quite rich. Some server platforms may provide access control functionality that goes beyond the UNIX-style mode attribute, but which is not as rich as the NFS ACL model. So that users can take advantage of this more limited functionality, the server may indicate that it supports ACLs as long as it follows the guidelines for mapping between its ACL model and the NFS version 4 ACL model.
The situation is complicated by the fact that a server may have multiple modules that enforce ACLs. For example, the enforcement for NFS version 4 access may be different from the enforcement for local access, and both may be different from the enforcement for access through other protocols such as SMB. So it may be useful for a server to accept an ACL even if not all of its modules are able to support it.
The guiding principle in all cases is that the server must not accept ACLs that appear to make the file more secure than it really is.
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The constants used for the type field (acetype4) are as follows:
const ACE4_ACCESS_ALLOWED_ACE_TYPE = 0x00000000;
const ACE4_ACCESS_DENIED_ACE_TYPE = 0x00000001;
const ACE4_SYSTEM_AUDIT_ACE_TYPE = 0x00000002;
const ACE4_SYSTEM_ALARM_ACE_TYPE = 0x00000003;
| Value | Abbreviation | Description |
|---|---|---|
| ACE4_ACCESS_ALLOWED_ACE_TYPE | ALLOW | Explicitly grants the access defined in acemask4 to the file or directory. |
| ACE4_ACCESS_DENIED_ACE_TYPE | DENY | Explicitly denies the access defined in acemask4 to the file or directory. |
| ACE4_SYSTEM_AUDIT_ACE_TYPE | AUDIT | LOG (system dependent) any access attempt to a file or directory which uses any of the access methods specified in acemask4. |
| ACE4_SYSTEM_ALARM_ACE_TYPE | ALARM | Generate a system ALARM (system dependent) when any access attempt is made to a file or directory for the access methods specified in acemask4. |
The "Abbreviation" column denotes how the types will be referred to throughout the rest of this document.
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A server need not support all of the above ACE types. The bitmask constants used to represent the above definitions within the aclsupport attribute are as follows:
const ACL4_SUPPORT_ALLOW_ACL = 0x00000001;
const ACL4_SUPPORT_DENY_ACL = 0x00000002;
const ACL4_SUPPORT_AUDIT_ACL = 0x00000004;
const ACL4_SUPPORT_ALARM_ACL = 0x00000008;
Clients should not attempt to set an ACE unless the server claims support for that ACE type. If the server receives a request to set an ACE that it cannot store, it MUST reject the request with NFS4ERR_ATTRNOTSUPP. If the server receives a request to set an ACE that it can store but cannot enforce, the server SHOULD reject the request with NFS4ERR_ATTRNOTSUPP.
Example: suppose a server can enforce NFS ACLs for NFS access but cannot enforce ACLs for local access. If arbitrary processes can run on the server, then the server SHOULD NOT indicate ACL support. On the other hand, if only trusted administrative programs run locally, then the server may indicate ACL support.
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The bitmask constants used for the access mask field are as follows:
const ACE4_READ_DATA = 0x00000001;
const ACE4_LIST_DIRECTORY = 0x00000001;
const ACE4_WRITE_DATA = 0x00000002;
const ACE4_ADD_FILE = 0x00000002;
const ACE4_APPEND_DATA = 0x00000004;
const ACE4_ADD_SUBDIRECTORY = 0x00000004;
const ACE4_READ_NAMED_ATTRS = 0x00000008;
const ACE4_WRITE_NAMED_ATTRS = 0x00000010;
const ACE4_EXECUTE = 0x00000020;
const ACE4_DELETE_CHILD = 0x00000040;
const ACE4_READ_ATTRIBUTES = 0x00000080;
const ACE4_WRITE_ATTRIBUTES = 0x00000100;
const ACE4_WRITE_RETENTION = 0x00000200;
const ACE4_WRITE_RETENTION_HOLD = 0x00000400;
const ACE4_DELETE = 0x00010000;
const ACE4_READ_ACL = 0x00020000;
const ACE4_WRITE_ACL = 0x00040000;
const ACE4_WRITE_OWNER = 0x00080000;
const ACE4_SYNCHRONIZE = 0x00100000;
| TOC |
ACE4_READ_DATA
Operation(s) affected:
READ
OPEN
Discussion:
Permission to read the data of the file.
Servers SHOULD allow a user the ability to read the data
of the file when only the ACE4_EXECUTE access mask bit is
allowed.
ACE4_LIST_DIRECTORY
Operation(s) affected:
READDIR
Discussion:
Permission to list the contents of a directory.
ACE4_WRITE_DATA
Operation(s) affected:
WRITE
OPEN
SETATTR of size
Discussion:
Permission to modify a file's data anywhere in the file's
offset range. This includes the ability to write to any
arbitrary offset and as a result to grow the file.
ACE4_ADD_FILE
Operation(s) affected:
CREATE
OPEN
Discussion:
Permission to add a new file in a directory. The CREATE
operation is affected when nfs_ftype4 is NF4LNK, NF4BLK,
NF4CHR, NF4SOCK, or NF4FIFO. (NF4DIR is not listed because
it is covered by ACE4_ADD_SUBDIRECTORY.) OPEN is affected
when used to create a regular file.
ACE4_APPEND_DATA
Operation(s) affected:
WRITE
OPEN
SETATTR of size
Discussion:
The ability to modify a file's data, but only starting at
EOF. This allows for the notion of append-only files, by
allowing ACE4_APPEND_DATA and denying ACE4_WRITE_DATA to
the same user or group. If a file has an ACL such as the
one described above and a WRITE request is made for
somewhere other than EOF, the server SHOULD return
NFS4ERR_ACCESS.
ACE4_ADD_SUBDIRECTORY
Operation(s) affected:
CREATE
Discussion:
Permission to create a subdirectory in a directory. The
CREATE operation is affected when nfs_ftype4 is NF4DIR.
ACE4_READ_NAMED_ATTRS
Operation(s) affected:
OPENATTR
Discussion:
Permission to read the named attributes of a file or to
lookup the named attributes directory. OPENATTR is
affected when it is not used to create a named attribute
directory. This is when 1.) createdir is TRUE, but a
named attribute directory already exists, or 2.) createdir
is FALSE.
ACE4_WRITE_NAMED_ATTRS
Operation(s) affected:
OPENATTR
Discussion:
Permission to write the named attributes of a file or
to create a named attribute directory. OPENATTR is
affected when it is used to create a named attribute
directory. This is when createdir is TRUE and no named
attribute directory exists. The ability to check whether
or not a named attribute directory exists depends on the
ability to look it up, therefore, users also need the
ACE4_READ_NAMED_ATTRS permission in order to create a
named attribute directory.
ACE4_EXECUTE
Operation(s) affected:
LOOKUP
READ
OPEN
Discussion:
Permission to execute a file or traverse/search a
directory.
Servers SHOULD allow a user the ability to read the data
of the file when only the ACE4_EXECUTE access mask bit is
allowed. This is because there is no way to execute a
file without reading the contents. Though a server may
treat ACE4_EXECUTE and ACE4_READ_DATA bits identically
when deciding to permit a READ operation, it SHOULD still
allow the two bits to be set independently in ACLs, and
MUST distinguish between them when replying to ACCESS
operations. In particular, servers SHOULD NOT silently
turn on one of the two bits when the other is set, as
that would make it impossible for the client to correctly
enforce the distinction between read and execute
permissions.
As an example, following a SETATTR of the following ACL:
nfsuser:ACE4_EXECUTE:ALLOW
A subsequent GETATTR of ACL for that file SHOULD return:
nfsuser:ACE4_EXECUTE:ALLOW
Rather than:
nfsuser:ACE4_EXECUTE/ACE4_READ_DATA:ALLOW
ACE4_DELETE_CHILD
Operation(s) affected:
REMOVE
Discussion:
Permission to delete a file or directory within a
directory. See section "ACE4_DELETE vs. ACE4_DELETE_CHILD"
for information on how these two access mask bits interact.
ACE4_READ_ATTRIBUTES
Operation(s) affected:
GETATTR of file system object attributes
Discussion:
The ability to read basic attributes (non-ACLs) of a file.
On a UNIX system, basic attributes can be thought of as
the stat level attributes. Allowing this access mask bit
would mean the entity can execute "ls -l" and stat.
ACE4_WRITE_ATTRIBUTES
Operation(s) affected:
SETATTR of time_access_set, time_backup,
time_create, time_modify_set, mimetype, hidden, system
Discussion:
Permission to change the times associated with a file
or directory to an arbitrary value. Also permission
to change the mimetype, hidden and system attributes.
A user having ACE4_WRITE_DATA permission, but lacking
ACE4_WRITE_ATTRIBUTES must be allowed to implicitly set
the times associated with a file.
ACE4_WRITE_RETENTION
Operation(s) affected:
SETATTR of retention_set, retentevt_set.
Discussion:
Permission to modify the durations of event and non-event-based
retention. Also permission to enable event and non-event-based
retention. A server MAY map ACE4_WRITE_ATTRIBUTES to
ACE_WRITE_RETENTION.
ACE4_WRITE_RETENTION_HOLD
Operation(s) affected:
SETATTR of retention_hold.
Discussion:
Permission to modify the administration retention holds.
A server MAY map ACE4_WRITE_ATTRIBUTES to
ACE_WRITE_RETENTION_HOLD.
ACE4_DELETE
Operation(s) affected:
REMOVE
Discussion:
Permission to delete the file or directory. See section
"ACE4_DELETE vs. ACE4_DELETE_CHILD" for information on how
these two access mask bits interact.
ACE4_READ_ACL
Operation(s) affected:
GETATTR of acl
Discussion:
Permission to read the ACL.
ACE4_WRITE_ACL
Operation(s) affected:
SETATTR of acl and mode
Discussion:
Permission to write the acl and mode attributes.
ACE4_WRITE_OWNER
Operation(s) affected:
SETATTR of owner and owner_group
Discussions:
Permission to write the owner and owner_group attributes.
On UNIX systems, this is the ability to execute chown().
ACE4_SYNCHRONIZE
Operation(s) affected:
NONE
Discussion:
Permission to access file locally at the server with
synchronized reads and writes.
Server implementations need not provide the granularity of control that is implied by this list of masks. For example, POSIX-based systems might not distinguish ACE4_APPEND_DATA (the ability to append to a file) from ACE4_WRITE_DATA (the ability to modify existing contents); both masks would be tied to a single "write" permission. When such a server returns attributes to the client, it would show both ACE4_APPEND_DATA and ACE4_WRITE_DATA if and only if the write permission is enabled.
If a server receives a SETATTR request that it cannot accurately implement, it should error in the direction of more restricted access. For example, suppose a server cannot distinguish overwriting data from appending new data, as described in the previous paragraph. If a client submits an ACE where ACE4_APPEND_DATA is set but ACE4_WRITE_DATA is not (or vice versa), the server should reject the request with NFS4ERR_ATTRNOTSUPP. Nonetheless, if the ACE has type DENY, the server may silently turn on the other bit, so that both ACE4_APPEND_DATA and ACE4_WRITE_DATA are denied.
| TOC |
Two access mask bits govern the ability to delete a file or directory object: ACE4_DELETE on the object itself, and ACE4_DELETE_CHILD on the object's parent directory.
Many systems also consult the "sticky bit" (MODE4_SVTX) and write mode bit on the parent directory when determining whether to allow a file to be deleted. The mode bit for write corresponds to ACE4_WRITE_DATA, which is the same physical bit as ACE4_ADD_FILE. Therefore, ACE4_ADD_FILE can come into play when determining permission to delete.
In the algorithm below, the strategy is that ACE4_DELETE and ACE4_DELETE_CHILD take precedence over the sticky bit, and the sticky bit takes precedence over the "write" mode bits (reflected in ACE4_ADD_FILE).
Server implementations SHOULD grant or deny permission to delete based on the following algorithm.
if ACE4_EXECUTE is denied by the parent directory ACL:
deny delete
else if ACE4_DELETE is allowed by the target object ACL:
allow delete
else if ACE4_DELETE_CHILD is allowed by the parent
directory ACL:
allow delete
else if ACE4_DELETE_CHILD is denied by the
parent directory ACL:
deny delete
else if ACE4_ADD_FILE is allowed by the parent directory ACL:
if MODE4_SVTX is set for the parent directory:
if the principal owns the parent directory OR
the principal owns the target object OR
ACE4_WRITE_DATA is allowed by the target
object ACL:
allow delete
else:
deny delete
else:
allow delete
else:
deny delete
| TOC |
The bitmask constants used for the flag field are as follows:
const ACE4_FILE_INHERIT_ACE = 0x00000001;
const ACE4_DIRECTORY_INHERIT_ACE = 0x00000002;
const ACE4_NO_PROPAGATE_INHERIT_ACE = 0x00000004;
const ACE4_INHERIT_ONLY_ACE = 0x00000008;
const ACE4_SUCCESSFUL_ACCESS_ACE_FLAG = 0x00000010;
const ACE4_FAILED_ACCESS_ACE_FLAG = 0x00000020;
const ACE4_IDENTIFIER_GROUP = 0x00000040;
A server need not support any of these flags. If the server supports flags that are similar to, but not exactly the same as, these flags, the implementation may define a mapping between the protocol-defined flags and the implementation-defined flags. Again, the guiding principle is that the file not appear to be more secure than it really is.
For example, suppose a client tries to set an ACE with ACE4_FILE_INHERIT_ACE set but not ACE4_DIRECTORY_INHERIT_ACE. If the server does not support any form of ACL inheritance, the server should reject the request with NFS4ERR_ATTRNOTSUPP. If the server supports a single "inherit ACE" flag that applies to both files and directories, the server may reject the request (i.e., requiring the client to set both the file and directory inheritance flags). The server may also accept the request and silently turn on the ACE4_DIRECTORY_INHERIT_ACE flag.
| TOC |
- ACE4_FILE_INHERIT_ACE
- Can be placed on a directory and indicates that this ACE should be added to each new non-directory file created.
- ACE4_DIRECTORY_INHERIT_ACE
- Can be placed on a directory and indicates that this ACE should be added to each new directory created.
- ACE4_INHERIT_ONLY_ACE
- Can be placed on a directory but does not apply to the directory; ALLOW and DENY ACEs with this bit set do not affect access to the directory, and AUDIT and ALARM ACEs with this bit set do not trigger log or alarm events. Such ACEs only take effect once they are applied (with this bit cleared) to newly created files and directories as specified by the above two flags.
- ACE4_NO_PROPAGATE_INHERIT_ACE
- Can be placed on a directory. This flag tells the server that inheritance of this ACE should stop at newly created child directories.
- ACE4_SUCCESSFUL_ACCESS_ACE_FLAG
- ACE4_FAILED_ACCESS_ACE_FLAG
- The ACE4_SUCCESSFUL_ACCESS_ACE_FLAG (SUCCESS) and ACE4_FAILED_ACCESS_ACE_FLAG (FAILED) flag bits relate only to ACE4_SYSTEM_AUDIT_ACE_TYPE (AUDIT) and ACE4_SYSTEM_ALARM_ACE_TYPE (ALARM) ACE types. If during the processing of the file's ACL, the server encounters an AUDIT or ALARM ACE that matches the principal attempting the OPEN, the server notes that fact, and the presence, if any, of the SUCCESS and FAILED flags encountered in the AUDIT or ALARM ACE. Once the server completes the ACL processing, it then notes if the operation succeeded or failed. If the operation succeeded, and if the SUCCESS flag was set for a matching AUDIT or ALARM ACE, then the appropriate AUDIT or ALARM event occurs. If the operation failed, and if the FAILED flag was set for the matching AUDIT or ALARM ACE, then the appropriate AUDIT or ALARM event occurs. Either or both of the SUCCESS or FAILED can be set, but if neither is set, the AUDIT or ALARM ACE is not useful.
- The previously described processing applies to that of the ACCESS operation as well, the difference being that "success" or "failure" does not mean whether ACCESS returns NFS4_OK or not. Success means whether ACCESS returns all requested and supported bits. Failure means whether ACCESS failed to return a bit that was requested and supported.
- ACE4_IDENTIFIER_GROUP
- Indicates that the "who" refers to a GROUP as defined under UNIX or a GROUP ACCOUNT as defined under Windows. Clients and servers must ignore the ACE4_IDENTIFIER_GROUP flag on ACEs with a who value equal to one of the special identifiers outlined in Section 6.2.1.5 (ACE Who).
| TOC |
The "who" field of an ACE is an identifier that specifies the principal or principals to whom the ACE applies. It may refer to a user or a group, with the flag bit ACE4_IDENTIFIER_GROUP specifying which.
There are several special identifiers which need to be understood universally, rather than in the context of a particular DNS domain. Some of these identifiers cannot be understood when an NFS client accesses the server, but have meaning when a local process accesses the file. The ability to display and modify these permissions is permitted over NFS, even if none of the access methods on the server understands the identifiers.
| Who | Description |
|---|---|
| OWNER | The owner of the file |
| GROUP | The group associated with the file. |
| EVERYONE | The world, including the owner and owning group. |
| INTERACTIVE | Accessed from an interactive terminal. |
| NETWORK | Accessed via the network. |
| DIALUP | Accessed as a dialup user to the server. |
| BATCH | Accessed from a batch job. |
| ANONYMOUS | Accessed without any authentication. |
| AUTHENTICATED | Any authenticated user (opposite of ANONYMOUS) |
| SERVICE | Access from a system service. |
| Table 7 |
To avoid conflict, these special identifiers are distinguish by an appended "@" and should appear in the form "xxxx@" (note: no domain name after the "@"). For example: ANONYMOUS@.
| TOC |
It is important to note that "EVERYONE@" is not equivalent to the UNIX "other" entity. This is because, by definition, UNIX "other" does not include the owner or owning group of a file. "EVERYONE@" means literally everyone, including the owner or owning group.
| TOC |
The NFS version 4 mode attribute is based on the UNIX mode bits. The following bits are defined:
const MODE4_SUID = 0x800; /* set user id on execution */
const MODE4_SGID = 0x400; /* set group id on execution */
const MODE4_SVTX = 0x200; /* save text even after use */
const MODE4_RUSR = 0x100; /* read permission: owner */
const MODE4_WUSR = 0x080; /* write permission: owner */
const MODE4_XUSR = 0x040; /* execute permission: owner */
const MODE4_RGRP = 0x020; /* read permission: group */
const MODE4_WGRP = 0x010; /* write permission: group */
const MODE4_XGRP = 0x008; /* execute permission: group */
const MODE4_ROTH = 0x004; /* read permission: other */
const MODE4_WOTH = 0x002; /* write permission: other */
const MODE4_XOTH = 0x001; /* execute permission: other */
Bits MODE4_RUSR, MODE4_WUSR, and MODE4_XUSR apply to the principal identified in the owner attribute. Bits MODE4_RGRP, MODE4_WGRP, and MODE4_XGRP apply to principals identified in the owner_group attribute but who are not identified in the owner attribute. Bits MODE4_ROTH, MODE4_WOTH, MODE4_XOTH apply to any principal that does not match that in the owner attribute, and does not have a group matching that of the owner_group attribute.
The remaining bits are not defined by this protocol. A server MUST NOT return bits other than those defined above in a GETATTR or READDIR operation, and it MUST return NFS4ERR_INVAL if bits other than those defined above are set in a SETATTR, CREATE, or OPEN operation.
| TOC |
The requirements in this section will be referred to in future sections, especially Section 6.4 (Requirements).
| TOC |
| TOC |
The server uses the algorithm described in Section 6.2.1 (ACL Attribute) to determine whether an ACL allows access to an object. However, the ACL may not be the sole determiner of access. For example:
| TOC |
Clients SHOULD NOT do their own access checks based on their interpretation the ACL, but rather use the OPEN and ACCESS operations to do access checks. This allows the client to act on the results of having the server determine whether or not access should be granted based on its interpretation of the ACL.
Clients must be aware of situations in which an object's ACL will define a certain access even though the server will not enforce it. In general, but especially in these situations, the client needs to do its part in the enforcement of access as defined by the ACL. To do this, the client MAY issue the appropriate ACCESS operation prior to servicing the request of the user or application in order to determine whether the user or application should be granted the access requested. For examples in which the ACL may define accesses that the server doesn't enforce see Section 6.3.1.1 (Server Considerations).
| TOC |
The following method can be used to calculate the MODE4_R*, MODE4_W* and MODE4_X* bits of a mode attribute, based upon an ACL.
| TOC |
The nine low-order mode bits (MODE4_R*, MODE4_W*, MODE4_X*) correspond to ACE4_READ_DATA, ACE4_WRITE_DATA/ACE4_APPEND_DATA, and ACE4_EXECUTE for OWNER@, GROUP@, and EVERYONE@. On some implementations, mode bits may represent a superset of these permissions, e.g. if a specific user is granted ACE4_WRITE_DATA, then MODE4_WGRP will be set, even though the file's owner_group is not granted ACE4_WRITE_DATA.
Server implementations are discouraged from doing this, as experience has shown that this is confusing and annoying to end users. The specifications above also discourage this practice to enforce the semantic that setting the mode attribute effectively specifies read, write, and execute for owner, group, and other.
| TOC |
The server that supports both mode and ACL must take care to synchronize the MODE4_*USR, MODE4_*GRP, and MODE4_*OTH bits with the ACEs which have respective who fields of "OWNER@", "GROUP@", and "EVERYONE@" so that the client can see semantically equivalent access permissions exist whether the client asks for owner, owner_group and mode attributes, or for just the ACL.
In this section, much is made of the methods in Section 6.3.2 (Computing a Mode Attribute from an ACL). Many requirements refer to this section. But note that the methods have behaviors specified with "SHOULD". This is intentional, to avoid invalidating existing implementations that compute the mode according to the withdrawn POSIX ACL draft (1003.1e draft 17), rather than by actual permissions on owner, group, and other.
| TOC |
| TOC |
When setting a mode attribute and not an ACL attribute, the mode attribute MUST be set as given. The ACL attribute MUST be modified such that the mode computed via the method in Section 6.3.2 (Computing a Mode Attribute from an ACL) yields the low-order nine bits (MODE4_R*, MODE4_W*, MODE4_X*) of the newly set mode attribute. The ACL SHOULD also be modified such that:
Access mask bits other those listed above, appearing in ALLOW ACEs, MAY also be disabled.
Note that ACEs with the flag ACE4_INHERIT_ONLY_ACE set do not affect the permissions of the ACL itself, nor do ACEs of the type AUDIT and ALARM. As such, it is desirable to leave these ACEs unmodified when modifying the ACL attribute.
Also note that the requirement may be met by discarding the ACL, in favor of an ACL that represents the mode and only the mode. This is permitted, but it is preferable for a server to preserve as much of the ACL as possible without violating the above requirements. Discarding the ACL makes it effectively impossible for a file created with a mode attribute to inherit an ACL (see Section 6.4.3 (Creating New Objects)).
| TOC |
When setting an ACL attribute and not a mode attribute, the ACL attribute SHOULD be set as given. The nine low-order bits of the mode attribute (MODE4_R*, MODE4_W*, MODE4_X*) MUST be modified to match the result of the method Section 6.3.2 (Computing a Mode Attribute from an ACL). The three high-order bits of the mode (MODE4_SUID, MODE4_SGID, MODE4_SVTX) SHOULD remain unchanged.
| TOC |
When setting both the mode and the ACL attribute in the same operation, the attributes MUST be applied in this order: mode, then ACL. The mode attribute is set as given, then the ACL attribute is set as given, possibly changing the final mode, as described above in Section 6.4.1.2 (Setting ACL and not mode).
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This section applies only to servers that support both the mode and the ACL attribute.
Some server implementations may have a concept of "objects without ACLs", meaning that all permissions are granted and denied according to the mode attribute, and that no ACL attribute is stored for that object. If an ACL attribute is requested of such a server, the server SHOULD return an ACL that does not conflict with the mode; that is to say, the ACL returned SHOULD represent the nine low-order bits of the mode attribute (MODE4_R*, MODE4_W*, MODE4_X*) as described in Section 6.3.2 (Computing a Mode Attribute from an ACL).
For other server implementations, the ACL attribute is always present for every object. Such servers SHOULD store at least the three high-order bits of the mode attribute (MODE4_SUID, MODE4_SGID, MODE4_SVTX). The server SHOULD return a mode attribute if one is requested, and the low-order nine bits of the mode (MODE4_R*, MODE4_W*, MODE4_X*) MUST match the result of applying the method in Section 6.3.2 (Computing a Mode Attribute from an ACL) to the ACL attribute.
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If a server supports the ACL attribute, it may use the ACL attribute on the parent directory to compute an initial ACL attribute for a newly created object. This will be referred to as the inherited ACL within this section. The act of adding one or more ACEs to the inherited ACL that are based upon ACEs in the parent directory's ACL will be referred to as inheriting an ACE within this section.
Implementors should standardize on what the behavior of CREATE and OPEN must be depending on the presence or absence of the mode and ACL attributes.
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If the object being created is not a directory, the inherited ACL SHOULD NOT inherit ACEs from the parent directory ACL unless the ACE4_FILE_INHERIT_FLAG is set.
If the object being created is a directory, the inherited ACL should inherit all inheritable ACEs from the parent directory, those that have ACE4_FILE_INHERIT_ACE or ACE4_DIRECTORY_INHERIT_ACE flag set. If the inheritable ACE has ACE4_FILE_INHERIT_ACE set, but ACE4_DIRECTORY_INHERIT_ACE is clear, the inherited ACE on the newly created directory MUST have the ACE4_INHERIT_ONLY_ACE flag set to prevent the directory from being affected by ACEs meant for non-directories.
If when a new directory is created and it inherits ACEs from its parent, for each inheritable ACE which affects the directory's permissions, a server MAY create two ACEs on the directory being created; one effective and one which is only inheritable (i.e. has ACE4_INHERIT_ONLY_ACE flag set). This gives the user and the server, in the cases which it must mask certain permissions upon creation, the ability to modify the effective permissions without modifying the ACE which is to be inherited to the new directory's children.
When a newly created object is created with attributes, and those attributes contain an ACL attribute and/or a mode attribute, the server MUST apply those attributes to the newly created object, as described in Section 6.4.1 (Setting the mode and/or ACL Attributes).
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This chapter describes the NFSv4 single-server name space. Single-server namespaces may be presented directly to clients, or they may be used as a basis to form larger multi-server namespaces (e.g. site-wide or organization-wide) to be presented to clients, as described in Section 10 (Multi-server Name Space).
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On a UNIX server, the name space describes all the files reachable by pathnames under the root directory or "/". On a Windows NT server the name space constitutes all the files on disks named by mapped disk letters. NFS server administrators rarely make the entire server's file system name space available to NFS clients. More often portions of the name space are made available via an "export" feature. In previous versions of the NFS protocol, the root filehandle for each export is obtained through the MOUNT protocol; the client sends a string that identifies the export of name space and the server returns the root filehandle for it. The MOUNT protocol supports an EXPORTS procedure that will enumerate the server's exports.
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The NFS version 4 protocol provides a root filehandle that clients can use to obtain filehandles for the exports of a particular server, via a series of LOOKUP operations within a COMPOUND, to traverse a path. A common user experience is to use a graphical user interface (perhaps a file "Open" dialog window) to find a file via progressive browsing through a directory tree. The client must be able to move from one export to another export via single-component, progressive LOOKUP operations.
This style of browsing is not well supported by the NFS version 2 and 3 protocols. The client expects all LOOKUP operations to remain within a single server file system. For example, the device attribute will not change. This prevents a client from taking name space paths that span exports.
An automounter on the client can obtain a snapshot of the server's name space using the EXPORTS procedure of the MOUNT protocol. If it understands the server's pathname syntax, it can create an image of the server's name space on the client. The parts of the name space that are not exported by the server are filled in with a "pseudo file system" that allows the user to browse from one mounted file system to another. There is a drawback to this representation of the server's name space on the client: it is static. If the server administrator adds a new export the client will be unaware of it.
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NFS version 4 servers avoid this name space inconsistency by presenting all the exports for a given server within the framework of a single namespace, for that server. An NFS version 4 client uses LOOKUP and READDIR operations to browse seamlessly from one export to another. Portions of the server name space that are not exported are bridged via a "pseudo file system" that provides a view of exported directories only. A pseudo file system has a unique fsid and behaves like a normal, read only file system.
Based on the construction of the server's name space, it is possible that multiple pseudo file systems may exist. For example,
/a pseudo file system
/a/b real file system
/a/b/c pseudo file system
/a/b/c/d real file system
Each of the pseudo file systems are considered separate entities and therefore will have its own unique fsid.
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The DOS and Windows operating environments are sometimes described as having "multiple roots". File Systems are commonly represented as disk letters. MacOS represents file systems as top level names. NFS version 4 servers for these platforms can construct a pseudo file system above these root names so that disk letters or volume names are simply directory names in the pseudo root.
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The nature of the server's pseudo file system is that it is a logical representation of file system(s) available from the server. Therefore, the pseudo file system is most likely constructed dynamically when the server is first instantiated. It is expected that the pseudo file system may not have an on disk counterpart from which persistent filehandles could be constructed. Even though it is preferable that the server provide persistent filehandles for the pseudo file system, the NFS client should expect that pseudo file system filehandles are volatile. This can be confirmed by checking the associated "fh_expire_type" attribute for those filehandles in question. If the filehandles are volatile, the NFS client must be prepared to recover a filehandle value (e.g. with a series of LOOKUP operations) when receiving an error of NFS4ERR_FHEXPIRED.
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If the server's root file system is exported, one might conclude that a pseudo-file system is unneeded. This not necessarily so. Assume the following file systems on a server:
/ disk1 (exported)
/a disk2 (not exported)
/a/b disk3 (exported)
Because disk2 is not exported, disk3 cannot be reached with simple LOOKUPs. The server must bridge the gap with a pseudo-file system.
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The server file system environment may be constructed in such a way that one file system contains a directory which is 'covered' or mounted upon by a second file system. For example:
/a/b (file system 1)
/a/b/c/d (file system 2)
The pseudo file system for this server may be constructed to look like:
/ (place holder/not exported)
/a/b (file system 1)
/a/b/c/d (file system 2)
It is the server's responsibility to present the pseudo file system that is complete to the client. If the client sends a lookup request for the path "/a/b/c/d", the server's response is the filehandle of the file system "/a/b/c/d". In previous versions of the NFS protocol, the server would respond with the filehandle of directory "/a/b/c/d" within the file system "/a/b".
The NFS client will be able to determine if it crosses a server mount point by a change in the value of the "fsid" attribute.
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The application of the server's security policy needs to be carefully considered by the implementor. One may choose to limit the viewability of portions of the pseudo file system based on the server's perception of the client's ability to authenticate itself properly. However, with the support of multiple security mechanisms and the ability to negotiate the appropriate use of these mechanisms, the server is unable to properly determine if a client will be able to authenticate itself. If, based on its policies, the server chooses to limit the contents of the pseudo file system, the server may effectively hide file systems from a client that may otherwise have legitimate access.
As suggested practice, the server should apply the security policy of a shared resource in the server's namespace to the components of the resource's ancestors. For example:
/
/a/b
/a/b/c
The /a/b/c directory is a real file system and is the shared resource. The security policy for /a/b/c is Kerberos with integrity. The server should apply the same security policy to /, /a, and /a/b. This allows for the extension of the protection of the server's namespace to the ancestors of the real shared resource.
For the case of the use of multiple, disjoint security mechanisms in the server's resources, the security for a particular object in the server's namespace should be the union of all security mechanisms of all direct descendants.
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Integrating locking into the NFS protocol necessarily causes it to be stateful. With the inclusion of such features as share reservations, file and directory delegations, recallable layouts, and support for mandatory byte-range locking the protocol becomes substantially more dependent on state than the traditional combination of NFS and NLM [XNFS]. There are three components to making this state manageable:
In this model, the server owns the state information. The client requests changes in locks and the server responds with the changes made. Non-client-initiated changes in locking state are infrequent and the client receives prompt notification of them and can adjust his view of the locking state to reflect the server's changes.
To support Win32 share reservations it is necessary to provide operations which atomically OPEN or CREATE files. Having a separate share/unshare operation would not allow correct implementation of the Win32 OpenFile API. In order to correctly implement share semantics, the previous NFS protocol mechanisms used when a file is opened or created (LOOKUP, CREATE, ACCESS) need to be replaced. The NFS version 4.1 protocol defines OPEN operation which looks up or creates a file and establishes locking state on the server.
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It is assumed that manipulating a lock is rare when compared to READ and WRITE operations. It is also assumed that crashes and network partitions are relatively rare. Therefore it is important that the READ and WRITE operations have a lightweight mechanism to indicate if they possess a held lock. A lock request contains the heavyweight information required to establish a lock and uniquely define the lock owner.
The following sections describe the transition from the heavyweight information to the eventual lightweight stateid used for most client and server locking interactions.
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A client must establish a clientid (see Section 2.4 (Client Identifiers)) and then one or more sessionids (see Section 2.9 (Session)) before performing any operations to open, lock, or delegate a file object. The sessionid services as a shorthand referral to an NFSv4.1 client.
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When opening a file or requesting a byte-range lock, the client must specify an identifier which represents the owner of the requested lock. This identifier is in the form of a state-owner, represented in the protocol by a state_owner4, a variable-length opaque array which, when concatenated with the current clientid uniquely defines the owner of lock managed by the client. This may be a thread id, process id, or other unique value.
Owners of opens and owners of byte-range locks are separate entities and remain separate even if the same opaque arrays are used to designate owners of each. The protocol distinguishes between open-owners (represented by open_owner4 structures) and lock-owners (represented by lock_owner4 structures).
Each open is associated with a specific open-owner while each byte-range lock is associated with a lock-owner and an open-owner, the latter being the open-owner associated with the open file under which the LOCK operation was done. Delegations and layouts, on the other hand, are not associated with a specific owner but are associated the client as a whole.
When the server grants a lock of any type (including opens, byte-range locks, delegations, and layouts) it responds with a unique stateid, that represents a set of locks (often a single lock) for the same file, of the same type, and sharing the same ownership characteristics. Thus opens of the same file by different open-owners each have an identifying stateid. Similarly, each set of byte-range locks on a file owned by a specific lock-owner and gotten via an open for a specific open-owner, has its own identifying stateid. Delegations and layouts also have associated stateid's by which they may be referenced. The stateid is used as a shorthand reference to a lock or set of locks and given a stateid the client can determine the associated state-owner or state-owners (in the case of an open-owner/lock-owner pair) and the associated. Clients, however, must not assume any such mapping and must not use a stateid returned for a given filehandle and state-owner in the context of a different filehandle or a different state-owner.
The server is free to form the stateid in any manner that it chooses as long as it is able to recognize invalid and out-of-date stateids. Although the protocol XDR definition divides the stateid into 'seqid' and 'other' fields, for the purposes of minor version one, this distinction is not important and the server may use the available space as it chooses, with one exception.
The exception is that stateids whose 'other' field is either all zeros or all ones are reserved and may not be generated by the server. Clients may use the protocol-defined special stateid values for their defined purposes, but any use of stateid's in this reserved class that are not specially defined by the protocol MUST result in an NFS4ERR_BAD_STATED being returned.
Clients may not compare stateids associated with different filehandles, so that a server might use stateids with the same bit pattern for all opens with a given open-owner or for all sets of byte-range locks associated with a given lock-owner/open-owner pair. However, if it does so, it must recognize and reject any use of stateid when the current filehandle is such that no lock for that filehandle by that open owner (or lock-owner/open-owner pair) exists.
Stateid's must remain valid until either a client reboot or a sever reboot or until the client returns all of the locks associated with the stateid by means of an operation such as CLOSE or DELEGRETURN. If the locks are lost due to revocation the stateid remains usable until the client frees it by using FREE_STATEID. Stateid's associated with byte-range locks are an exception. They remain valid even if a LOCKU free all remaining locks, so long as the open file with which they are associated remains open, unless the client does a FREE_STATEID to caused the stateid to be freed.
Because each operation using a stateid occurs as part of a session, each stateid is implicitly associated with the clientid assigned to that session. Use of a stateid in the context of a session where the clientid is invalid should result in the error NFS4ERR_STALE_STATEID. Servers MUST NOT do any validation or return other errors in this case, even if they have sufficient information available to validate stateids associated with an out-of-date client.
One mechanism that may be used to satisfy the requirement that the server recognize invalid and out-of-date stateids is for the server to divide the stateid into two fields. This division may coincide with the documented division into 'seqid' and 'other' fields or it may divide the stateid field up in any other ay it chooses.
And then store in each table entry,
With this information, the following procedure would be used to validate an incoming stateid and return an appropriate error, when necessary:
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All READ, WRITE and SETATTR operations contain a stateid. For the purposes of this section, SETATTR operations which change the size attribute of a file are treated as if they are writing the area between the old and new size (i.e. the range truncated or added to the file by means of the SETATTR), even where SETATTR is not explicitly mentioned in the text.
If the state-owner performs a READ or WRITE in a situation in which it has established a lock or share reservation on the server (any OPEN constitutes a share reservation) the stateid (previously returned by the server) must be used to indicate what locks, including both record locks and share reservations, are held by the state-owner. If no state is established by the client, either record lock or share reservation, a special stateid of all bits 0 (including all fields of the stateid) is used. Regardless whether a stateid of all bits 0, or a stateid returned by the server is used, if there is a conflicting share reservation or mandatory record lock held on the file, the server MUST refuse to service the READ or WRITE operation.
Share reservations are established by OPEN operations and by their nature are mandatory in that when the OPEN denies READ or WRITE operations, that denial results in such operations being rejected with error NFS4ERR_LOCKED. Record locks may be implemented by the server as either mandatory or advisory, or the choice of mandatory or advisory behavior may be determined by the server on the basis of the file being accessed (for example, some UNIX-based servers support a "mandatory lock bit" on the mode attribute such that if set, record locks are required on the file before I/O is possible). When record locks are advisory, they only prevent the granting of conflicting lock requests and have no effect on READs or WRITEs. Mandatory record locks, however, prevent conflicting I/O operations. When they are attempted, they are rejected with NFS4ERR_LOCKED. When the client gets NFS4ERR_LOCKED on a file it knows it has the proper share reservation for, it will need to issue a LOCK request on the region of the file that includes the region the I/O was to be performed on, with an appropriate locktype (i.e. READ*_LT for a READ operation, WRITE*_LT for a WRITE operation).
Note that for UNIX environments that support mandatory file locking, the distinction between advisory and mandatory locking is subtle. In fact, advisory and mandatory record locks are exactly the same in so far as the APIs and requirements on implementation. If the mandatory lock attribute is set on the file, the server checks to see if the lock-owner has an appropriate shared (read) or exclusive (write) record lock on the region it wishes to read or write to. If there is no appropriate lock, the server checks if there is a conflicting lock (which can be done by attempting to acquire the conflicting lock on the behalf of the lock-owner, and if successful, release the lock after the READ or WRITE is done), and if there is, the server returns NFS4ERR_LOCKED.
For Windows environments, there are no advisory record locks, so the server always checks for record locks during I/O requests.
Thus, the NFS version 4 LOCK operation does not need to distinguish between advisory and mandatory record locks. It is the NFS version 4 server's processing of the READ and WRITE operations that introduces the distinction.
Every stateid other than the special stateid values noted in this section, whether returned by an OPEN-type operation (i.e. OPEN, OPEN_DOWNGRADE), or by a LOCK-type operation (i.e. LOCK or LOCKU), defines an access mode for the file (i.e. READ, WRITE, or READ-WRITE) as established by the original OPEN which caused the allocation of the open stateid and as modified by subsequent OPENs and OPEN_DOWNGRADEs for the same open-owner/file pair. Stateids returned by byte-range lock operations imply the access mode for the open stateid associated with the lock set represented by the stateid. Delegation stateids have an access mode based on the type of delegation. When a READ, WRITE, or SETATTR which specifies the size attribute, is done, the operation is subject to checking against the access mode to verify that the operation is appropriate given the OPEN with which the operation is associated.
In the case of WRITE-type operations (i.e. WRITEs and SETATTRs which set size), the server must verify that the access mode allows writing and return an NFS4ERR_OPENMODE error if it does not. In the case, of READ, the server may perform the corresponding check on the access mode, or it may choose to allow READ on opens for WRITE only, to accommodate clients whose write implementation may unavoidably do reads (e.g. due to buffer cache constraints). However, even if READs are allowed in these circumstances, the server MUST still check for locks that conflict with the READ (e.g. another open specify denial of READs). Note that a server which does enforce the access mode check on READs need not explicitly check for conflicting share reservations since the existence of OPEN for read access guarantees that no conflicting share reservation can exist.
A special stateid of all bits 1 (one), including all fields in the stateid indicates a desire to bypass locking checks. The server MAY allow READ operations to bypass locking checks at the server, when this special stateid is used. However, WRITE operations with this special stateid value MUST NOT bypass locking checks and are treated exactly the same as if a stateid of all bits 0 were used.
A lock may not be granted while a READ or WRITE operation using one of the special stateids is being performed and the range of the lock request conflicts with the range of the READ or WRITE operation. For the purposes of this paragraph, a conflict occurs when a shared lock is requested and a WRITE operation is being performed, or an exclusive lock is requested and either a READ or a WRITE operation is being performed. A SETATTR that sets size is treated similarly to a WRITE as discussed above.
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The protocol allows a lock owner to request a lock with a byte range and then either upgrade, downgrade, or unlock a sub-range of the initial lock. It is expected that this will be an uncommon type of request. In any case, servers or server filesystems may not be able to support sub-range lock semantics. In the event that a server receives a locking request that represents a sub-range of current locking state for the lock owner, the server is allowed to return the error NFS4ERR_LOCK_RANGE to signify that it does not support sub-range lock operations. Therefore, the client should be prepared to receive this error and, if appropriate, report the error to the requesting application.
The client is discouraged from combining multiple independent locking ranges that happen to be adjacent into a single request since the server may not support sub-range requests and for reasons related to the recovery of file locking state in the event of server failure. As discussed in the section "Server Failure and Recovery" below, the server may employ certain optimizations during recovery that work effectively only when the client's behavior during lock recovery is similar to the client's locking behavior prior to server failure.
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If a client has a write lock on a record, it can request an atomic downgrade of the lock to a read lock via the LOCK request, by setting the type to READ_LT. If the server supports atomic downgrade, the request will succeed. If not, it will return NFS4ERR_LOCK_NOTSUPP. The client should be prepared to receive this error, and if appropriate, report the error to the requesting application.
If a client has a read lock on a record, it can request an atomic upgrade of the lock to a write lock via the LOCK request by setting the type to WRITE_LT or WRITEW_LT. If the server does not support atomic upgrade, it will return NFS4ERR_LOCK_NOTSUPP. If the upgrade can be achieved without an existing conflict, the request will succeed. Otherwise, the server will return either NFS4ERR_DENIED or NFS4ERR_DEADLOCK. The error NFS4ERR_DEADLOCK is returned if the client issued the LOCK request with the type set to WRITEW_LT and the server has detected a deadlock. The client should be prepared to receive such errors and if appropriate, report the error to the requesting application.
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Some clients require the support of blocking locks. NFSv4.1 does not provide a callback when a previously unavailable lock becomes available. Clients thus have no choice but to continually poll for the lock. This presents a fairness problem. Two new lock types are added, READW and WRITEW, and are used to indicate to the server that the client is requesting a blocking lock. The server should maintain an ordered list of pending blocking locks. When the conflicting lock is released, the server may wait the lease period for the first waiting client to re-request the lock. After the lease period expires the next waiting client request is allowed the lock. Clients are required to poll at an interval sufficiently small that it is likely to acquire the lock in a timely manner. The server is not required to maintain a list of pending blocked locks as it is used to increase fairness and not correct operation. Because of the unordered nature of crash recovery, storing of lock state to stable storage would be required to guarantee ordered granting of blocking locks.
Servers may also note the lock types and delay returning denial of the request to allow extra time for a conflicting lock to be released, allowing a successful return. In this way, clients can avoid the burden of needlessly frequent polling for blocking locks. The server should take care in the length of delay in the event the client retransmits the request.
If a server receives a blocking lock request, denies it, and then later receives a nonblocking request for the same lock, which is also denied, then it should remove the lock in question from its list of pending blocking locks. Clients should use such a nonblocking request to indicate to the server that this is the last time they intend to poll for the lock, as may happen when the process requesting the lock is interrupted. This is a courtesy to the server, to prevent it from unnecessarily waiting a lease period before granting other lock requests. However, clients are not required to perform this courtesy, and servers must not depend on them doing so. Also, clients must be prepared for the possibility that this final locking request will be accepted.
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The purpose of a lease is to allow a server to remove stale locks that are held by a client that has crashed or is otherwise unreachable. It is not a mechanism for cache consistency and lease renewals may not be denied if the lease interval has not expired.
Since each session is associated with a specific client, any operation issued on that session is an indication that the associated client is reachable. When a request is issued for a given session, execution of a SEQUENCE operation will result in all leases for the associated client to be implicitly renewed. This approach allows for low overhead lease renewal which scales well. In the typical case no extra RPC calls are required for lease renewal and in the worst case one RPC is required every lease period, via a COMPOUND that consists solely of a single SEQUENCE operation. The number of locks held by the client is not a factor since all state for the client is involved with the lease renewal action.
Since all operations that create a new lease also renew existing leases, the server must maintain a common lease expiration time for all valid leases for a given client. This lease time can then be easily updated upon implicit lease renewal actions.
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The important requirement in crash recovery is that both the client and the server know when the other has failed. Additionally, it is required that a client sees a consistent view of data across server restarts or reboots. All READ and WRITE operations that may have been queued within the client or network buffers must wait until the client has successfully recovered the locks protecting the READ and WRITE operations.
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In the event that a client fails, the server may release the client's locks when the associated leases have expired. Conflicting locks from another client may only be granted after this lease expiration. When a client has not failed and re-establishes his lease before expiration occurs, requests for conflicting locks will not be granted.
To minimize client delay upon restart, lock requests are associated with an instance of the client by a client supplied verifier. This verifier is part of the initial EXCHANGE_ID call made by the client. The server returns a clientid as a result of the EXCHANGE_ID operation. The client then confirms the use of the clientid by establishing a session associated with that clientid. All locks, including opens, byte-range locks, delegations, and layout obtained by sessions using that clientid are associated with that clientid.
Since the verifier will be changed by the client upon each initialization, the server can compare a new verifier to the verifier associated with currently held locks and determine that they do not match. This signifies the client's new instantiation and subsequent loss of locking state. As a result, the server is free to release all locks held which are associated with the old clientid which was derived from the old verifier. At this point conflicting locks from other clients, kept waiting while the leaser had not yet expired, can be granted.
Note that the verifier must have the same uniqueness properties of the verifier for the COMMIT operation.
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If the server loses locking state (usually as a result of a restart or reboot), it must allow clients time to discover this fact and re-establish the lost locking state. The client must be able to re-establish the locking state without having the server deny valid requests because the server has granted conflicting access to another client. Likewise, if there is a possibility that clients have not yet re-established their locking state for a file, the server must disallow READ and WRITE operations for that file.
A client can determine that server failure (and thus loss of locking state) has occurred, when it receives one of two errors. The NFS4ERR_STALE_STATEID error indicates a stateid invalidated by a reboot or restart. The NFS4ERR_STALE_CLIENTID error indicates a clientid invalidated by reboot or restart. When either of these are received, the client must establish a new clientid (See Section 8.1.1 (Client and Session ID)) and re-establish its locking state.
Once a session is established using the new clientid, the client will use reclaim-type locking requests (i.e. LOCK requests with reclaim set to true and OPEN operations with a claim type of CLAIM_PREVIOUS) to re-establish its locking state. Once this is done, or if there is no such locking state to reclaim, the client does a RECLAIM_COMPLETE operation to indicate that it has reclaimed all of the locking state that it will reclaim. Once a client does a RECLAIM_COMPLETE operation, it may attempt non-reclaim locking operations, although it may get NFS4ERR_GRACE errors on these until the period of special handling is over.
The period of special handling of locking and READs and WRITEs, is referred to as the "grace period". During the grace period, clients recover locks and the associated state using reclaim-type locking requests. During this period, the server must reject READ and WRITE operations and non-reclaim locking requests (i.e. other LOCK and OPEN operations) with an error of NFS4ERR_GRACE, unless it is able to guarantee that these may be done safely, as described below.
The grace period may last until all clients who are known to possibly have had locks have done a RECLAIM_COMPLETE operation, indicating that they have finished reclaiming the locks they held before the server reboot. The server is assumed to maintain in stable storage a list of clients who may have such locks. The server may also terminate the grace period before all clients have done RECLAIM_COMPLETE. The server SHOULD NOT terminate the grace period before a time equal to the lease period in order to give clients an opportunity to find out about the server reboot. Some additional time in order to allow time to establish a new clientid and session and to effect lock reclaims may be added.
If the server can reliably determine that granting a non-reclaim request will not conflict with reclamation of locks by other clients, the NFS4ERR_GRACE error does not have to be returned even within the grace period, although NFS4ERR_GRACE must always be returned to clients attempting a non-reclaim lock request before doing their own RECLAIM_COMPLETE. For the server to be able to service READ and WRITE operations during the grace period, it must again be able to guarantee that no possible conflict could arise between a potential reclaim locking request and the READ or WRITE operation. If the server is unable to offer that guarantee, the NFS4ERR_GRACE error must be returned to the client.
For a server to provide simple, valid handling during the grace period, the easiest method is to simply reject all non-reclaim locking requests and READ and WRITE operations by returning the NFS4ERR_GRACE error. However, a server may keep information about granted locks in stable storage. With this information, the server could determine if a regular lock or READ or WRITE operation can be safely processed.
For example, if the server maintained on stable storage summary information on whether mandatory locks exist, either mandatory byte-range locks, or share reservations specifying deny modes, many requests could be allowed during the grace period. If it is known that no such share reservations exist, OPEN request that do not specify deny modes may be safely granted. If, in addition, it is known that no mandatory byte-range locks exist, either through information stored on stable storage or simply because the server does not support such locks, READ and WRITE requests may be safely processed during the grace period.
To reiterate, for a server that allows non-reclaim lock and I/O requests to be processed during the grace period, it MUST determine that no lock subsequently reclaimed will be rejected and that no lock subsequently reclaimed would have prevented any I/O operation processed during the grace period.
Clients should be prepared for the return of NFS4ERR_GRACE errors for non-reclaim lock and I/O requests. In this case the client should employ a retry mechanism for the request. A delay (on the order of several seconds) between retries should be used to avoid overwhelming the server. Further discussion of the general issue is included in [Floyd]. The client must account for the server that is able to perform I/O and non-reclaim locking requests within the grace period as well as those that can not do so.
A reclaim-type locking request outside the server's grace period can only succeed if the server can guarantee that no conflicting lock or I/O request has been granted since reboot or restart.
A server may, upon restart, establish a new value for the lease period. Therefore, clients should, once a new clientid is established, refetch the lease_time attribute and use it as the basis for lease renewal for the lease associated with that server. However, the server must establish, for this restart event, a grace period at least as long as the lease period for the previous server instantiation. This allows the client state obtained during the previous server instance to be reliably re-established.
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If the duration of a network partition is greater than the lease period provided by the server, the server will have not received a lease renewal from the client. If this occurs, the server may free all locks held for the client, or it may allow the lock state to remain for a considerable period, subject to the constraint that if a request for a conflicting lock is made, locks associated with expired leases do not prevent such a conflicting lock from being granted but are revoked as necessary so as not to interfere with such conflicting requests.
If the server chooses to delay freeing of lock state until there is a conflict, it may either free all of the clients locks once there is a conflict, or it may only revoke the minimum set of locks necessary to allow conflicting requests. When it adopts the finer-grained approach, it must revoke all locks associated with a given stateid, as long as it revokes a single such lock.
When the server chooses to free all of a client's lock state, either immediately upon lease expiration, or a result of the first attempt to get a lock, all stateids held by the client will become invalid or stale. Once the client is able to reach the server after such a network partition, the status returned by the SEQUENCE operation will indicate a loss of locking state. In addition all I/O submitted by the client with the now invalid stateids will fail with the server returning the error NFS4ERR_EXPIRED. Once the client learns of the loss of locking state, it will suitably notify the applications that held the invalidated locks. The client should then take action to free invalidated stateid's, either by establishing a new client id using a new verifier or by doing a FREE_STATEID operation to release each of the invalidated stateid's.
When the server adopts a finer-grained approach to revocation of locks when lease have expired, only a subset of stateids will normally become invalid during a network partition. When the client is able to communicate with the server after such a network partition, the status returned by the SEQUENCE operation will indicate a partial loss of locking state. In addition, operations, including I/O submitted by the client with the now invalid stateids will fail with the server returning the error NFS4ERR_EXPIRED. Once the client learns of the loss of locking state, it will use the TEST_STATEID operation on all of its stateid's to determine which locks have been lost and them suitably notify the applications that held the invalidated locks. The client can then release the invalidated locking state and acknowledge the revocation of the associated locks by doing a FREE_STATEID operation on each of the invalidated stateid's.
When a network partition is combined with a server reboot, there are edge conditions that place requirements on the server in order to avoid silent data corruption following the server reboot. Two of these edge conditions are known, and are discussed below.
The first edge condition arises as a result of the scenarios such as the following:
Thus, at the final step, the server has erroneously granted client A's lock reclaim. If client B modified the object the lock was protecting, client A will experience object corruption.
The second known edge condition arises in situations such as the following:
As with the first edge condition, the final step of the scenario of the second edge condition has the server erroneously granting client A's lock reclaim.
Solving the first and second edge conditions requires that the server either always assumes after it reboots that some edge condition occurs, and thus return NFS4ERR_NO_GRACE for all reclaim attempts, or that the server record some information in stable storage. The amount of information the server records in stable storage is in inverse proportion to how harsh the server intends to be whenever edge conditions arise. The server that is completely tolerant of all edge conditions will record in stable storage every lock that is acquired, removing the lock record from stable storage only when the lock is released. For the two edge conditions discussed above, the harshest a server can be, and still support a grace period for reclaims, requires that the server record in stable storage information some minimal information. For example, a server implementation could, for each client, save in stable storage a record containing:
Assuming the above record keeping, for the first edge condition, after the server reboots, the record that client A's lease expired means that another client could have acquired a conflicting record lock, share reservation, or delegation. Hence the server must reject a reclaim from client A with the error NFS4ERR_NO_GRACE.
For the second edge condition, after the server reboots for a second time, the indication that the client had not completed its reclaims at the time at which the grace period ended means that the server must reject a reclaim from client A with the error NFS4ERR_NO_GRACE.
When either edge condition occurs, the client's attempt to reclaim locks will result in the error NFS4ERR_NO_GRACE. When this is received, or after the client reboots with no lock state, the client will issue a RECLAIM_COMPLETE. When the RECLAIM_COMPLETE is received, the server and client are again in agreement regarding reclaimable locks and both booleans in persistent storage can be reset, to be set again only when there is a subsequent event that causes lock reclaim operations to be questionable.
Regardless of the level and approach to record keeping, the server MUST implement one of the following strategies (which apply to reclaims of share reservations, record locks, and delegations):
A mandate for the client's handling of the NFS4ERR_NO_GRACE error is outside the scope of this specification, since the strategies for such handling are very dependent on the client's operating environment. However, one potential approach is described below.
When the client receives NFS4ERR_NO_GRACE, it could examine the change attribute of the objects the client is trying to reclaim state for, and use that to determine whether to re-establish the state via normal OPEN or LOCK requests. This is acceptable provided the client's operating environment allows it. In other words, the client implementor is advised to document for his users the behavior. The client could also inform the application that its record lock or share reservations (whether they were delegated or not) have been lost, such as via a UNIX signal, a GUI pop-up window, etc. See the section, "Data Caching and Revocation" for a discussion of what the client should do for dealing with unreclaimed delegations on client state.
For further discussion of revocation of locks see Section 8.7 (Server Revocation of Locks).
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At any point, the server can revoke locks held by a client and the client must be prepared for this event. When the client detects that its locks have been or may have been revoked, the client is responsible for validating the state information between itself and the server. Validating locking state for the client means that it must verify or reclaim state for each lock currently held.
The first occasion of lock revocation is upon server reboot or re-initialization. In this instance the client will receive an error (NFS4ERR_STALE_STATEID or NFS4ERR_STALE_CLIENTID) and the client will proceed with normal crash recovery as described in the previous section.
The second occasion of lock revocation is the inability to renew the lease before expiration, as discussed above. While this is considered a rare or unusual event, the client must be prepared to recover. The server is responsible for determining lease expiration, and deciding exactly how to deal with it, informing the client of the scope of the lock revocation. The client then uses the status information provided by the server to synchronize his locking state with that of the server, in order to recover.
The third occasion of lock revocation can occur as a result of revocation of locks within the lease period, either because of administrative intervention, or because a recallable lock (a delegation or layout) was not returned within the lease period after having been recalled. While these are considered rare events, they are possible and the client must be prepared to deal with them. When either of these events occur, the client finds out about the situation through the status returned by the SEQUENCE operation. Any use of stateids associated with revoked locks will receive the error NFS4ERR_ADMIN_REVOKED or NFS4ERR_DELEG_REVOKED, as appropriate.
In all situations in which a subset of locking state may have been revoked, which include all cases in which locking state is revoked within the lease period, it is up to the client to determine which locks have been revoked and which have not. It does this by using the TEST_STATEID operation on the appropriate set of stateid's. Once the set of revoked locks has been determined, the applications can be notified, and the invalidated stateid's can be freed and lock revocation acknowledged by using FREE_STATEID.
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A share reservation is a mechanism to control access to a file. It is a separate and independent mechanism from record locking. When a client opens a file, it issues an OPEN operation to the server specifying the type of access required (READ, WRITE, or BOTH) and the type of access to deny others (deny NONE, READ, WRITE, or BOTH). If the OPEN fails the client will fail the application's open request.
Pseudo-code definition of the semantics:
if (request.access == 0)
return (NFS4ERR_INVAL)
else
if ((request.access & file_state.deny)) ||
(request.deny & file_state.access))
return (NFS4ERR_DENIED)
This checking of share reservations on OPEN is done with no exception for an existing OPEN for the same open-owner.
The constants used for the OPEN and OPEN_DOWNGRADE operations for the access and deny fields are as follows:
const OPEN4_SHARE_ACCESS_READ = 0x00000001;
const OPEN4_SHARE_ACCESS_WRITE = 0x00000002;
const OPEN4_SHARE_ACCESS_BOTH = 0x00000003;
const OPEN4_SHARE_DENY_NONE = 0x00000000;
const OPEN4_SHARE_DENY_READ = 0x00000001;
const OPEN4_SHARE_DENY_WRITE = 0x00000002;
const OPEN4_SHARE_DENY_BOTH = 0x00000003;
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To provide correct share semantics, a client MUST use the OPEN operation to obtain the initial filehandle and indicate the desired access and what if any access to deny. Even if the client intends to use a stateid of all 0's or all 1's, it must still obtain the filehandle for the regular file with the OPEN operation so the appropriate share semantics can be applied. For clients that do not have a deny mode built into their open programming interfaces, deny equal to NONE should be used.
The OPEN operation with the CREATE flag, also subsumes the CREATE operation for regular files as used in previous versions of the NFS protocol. This allows a create with a share to be done atomically.
The CLOSE operation removes all share reservations held by the open-owner on that file. If record locks are held, the client SHOULD release all locks before issuing a CLOSE. The server MAY free all outstanding locks on CLOSE but some servers may not support the CLOSE of a file that still has record locks held. The server MUST return failure, NFS4ERR_LOCKS_HELD, if any locks would exist after the CLOSE.
The LOOKUP operation will return a filehandle without establishing any lock state on the server. Without a valid stateid, the server will assume the client has the least access. For example, a file opened with deny READ/WRITE cannot be accessed using a filehandle obtained through LOOKUP because it would not have a valid stateid (i.e. using a stateid of all bits 0 or all bits 1).
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When an OPEN is done for a file and the open-owner for which the open is being done already has the file open, the result is to upgrade the open file status maintained on the server to include the access and deny bits specified by the new OPEN as well as those for the existing OPEN. The result is that there is one open file, as far as the protocol is concerned, and it includes the union of the access and deny bits for all of the OPEN requests completed. Only a single CLOSE will be done to reset the effects of both OPENs. Note that the client, when issuing the OPEN, may not know that the same file is in fact being opened. The above only applies if both OPENs result in the OPENed object being designated by the same filehandle.
When the server chooses to export multiple filehandles corresponding to the same file object and returns different filehandles on two different OPENs of the same file object, the server MUST NOT "OR" together the access and deny bits and coalesce the two open files. Instead the server must maintain separate OPENs with separate stateids and will require separate CLOSEs to free them.
When multiple open files on the client are merged into a single open file object on the server, the close of one of the open files (on the client) may necessitate change of the access and deny status of the open file on the server. This is because the union of the access and deny bits for the remaining opens may be smaller (i.e. a proper subset) than previously. The OPEN_DOWNGRADE operation is used to make the necessary change and the client should use it to update the server so that share reservation requests by other clients are handled properly.
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When determining the time period for the server lease, the usual lease tradeoffs apply. Short leases are good for fast server recovery at a cost of increased operations to effect lease renewal (when there are no other operations during the period to effect lease renewal as a side-effect). Long leases are certainly kinder and gentler to servers trying to handle very large numbers of clients. The number of extra requests to effect lock renewal drop in inverse proportion to the lease time. The disadvantages of long leases include the possibility of slower recovery after certain failures. After server failure, a longer grace period may be required when some clients do not promptly reclaim their locks and do a RECLAIM_COMPLETE. In the event of client failure, it can longer period for leases to expire thus forcing conflicting requests to wait.
Long leases are usable if the server is able to store lease state in non-volatile memory. Upon recovery, the server can reconstruct the lease state from its non-volatile memory and continue operation with its clients and therefore long leases would not be an issue.
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To avoid the need for synchronized clocks, lease times are granted by the server as a time delta. However, there is a requirement that the client and server clocks do not drift excessively over the duration of the lock. There is also the issue of propagation delay across the network which could easily be several hundred milliseconds as well as the possibility that requests will be lost and need to be retransmitted.
To take propagation delay into account, the client should subtract it from lease times (e.g. if the client estimates the one-way propagation delay as 200 msec, then it can assume that the lease is already 200 msec old when it gets it). In addition, it will take another 200 msec to get a response back to the server. So the client must send a lock renewal or write data back to the server 400 msec before the lease would expire.
The server's lease period configuration should take into account the network distance of the clients that will be accessing the server's resources. It is expected that the lease period will take into account the network propagation delays and other network delay factors for the client population. Since the protocol does not allow for an automatic method to determine an appropriate lease period, the server's administrator may have to tune the lease period.
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There are a number of operations and fields within existing operations that no longer have a function in minor version one. In one way or another, these changes are all due to the implementation of sessions which provides client context and replay protection as a base feature of the protocol, separate from locking itself.
The following operations have become mandatory-to-not-implement. The server should return NFS4ERR_NOTSUPP if these operations are found in an NFSv4.1 COMPOUND.
Also, there are a number of fields, present in existing operations related to locking that have no use in minor version one. They were used in minor version zero to perform functions now provided in a different fashion.
Such vestigial fields in existing operations should be set by the client to zero. When they are not, the server MUST return an NFS4ERR_INVAL error.
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Client-side caching of data, of file attributes, and of file names is essential to providing good performance with the NFS protocol. Providing distributed cache coherence is a difficult problem and previous versions of the NFS protocol have not attempted it. Instead, several NFS client implementation techniques have been used to reduce the problems that a lack of coherence poses for users. These techniques have not been clearly defined by earlier protocol specifications and it is often unclear what is valid or invalid client behavior.
The NFS version 4 protocol uses many techniques similar to those that have been used in previous protocol versions. The NFS version 4 protocol does not provide distributed cache coherence. However, it defines a more limited set of caching guarantees to allow locks and share reservations to be used without destructive interference from client side caching.
In addition, the NFS version 4 protocol introduces a delegation mechanism which allows many decisions normally made by the server to be made locally by clients. This mechanism provides efficient support of the common cases where sharing is infrequent or where sharing is read-only.
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Caching techniques used in previous versions of the NFS protocol have been successful in providing good performance. However, several scalability challenges can arise when those techniques are used with very large numbers of clients. This is particularly true when clients are geographically distributed which classically increases the latency for cache revalidation requests.
The previous versions of the NFS protocol repeat their file data cache validation requests at the time the file is opened. This behavior can have serious performance drawbacks. A common case is one in which a file is only accessed by a single client. Therefore, sharing is infrequent.
In this case, repeated reference to the server to find that no conflicts exist is expensive. A better option with regards to performance is to allow a client that repeatedly opens a file to do so without reference to the server. This is done until potentially conflicting operations from another client actually occur.
A similar situation arises in connection with file locking. Sending file lock and unlock requests to the server as well as the read and write requests necessary to make data caching consistent with the locking semantics (see the section "Data Caching and File Locking") can severely limit performance. When locking is used to provide protection against infrequent conflicts, a large penalty is incurred. This penalty may discourage the use of file locking by applications.
The NFS version 4 protocol provides more aggressive caching strategies with the following design goals:
.IP o Compatibility with a large range of server semantics. .IP o Provide the same caching benefits as previous versions of the NFS protocol when unable to provide the more aggressive model. .IP o Requirements for aggressive caching are organized so that a large portion of the benefit can be obtained even when not all of the requirements can be met. .LP The appropriate requirements for the server are discussed in later sections in which specific forms of caching are covered. (see the section "Open Delegation").
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Recallable delegation of server responsibilities for a file to a client improves performance by avoiding repeated requests to the server in the absence of inter-client conflict. With the use of a "callback" RPC from server to client, a server recalls delegated responsibilities when another client engages in sharing of a delegated file.
A delegation is passed from the server to the client, specifying the object of the delegation and the type of delegation. There are different types of delegations but each type contains a stateid to be used to represent the delegation when performing operations that depend on the delegation. This stateid is similar to those associated with locks and share reservations but differs in that the stateid for a delegation is associated with a clientid and may be used on behalf of all the open_owners for the given client. A delegation is made to the client as a whole and not to any specific process or thread of control within it.
Because callback RPCs may not work in all environments (due to firewalls, for example), correct protocol operation does not depend on them. Preliminary testing of callback functionality by means of a CB_NULL procedure determines whether callbacks can be supported. The CB_NULL procedure checks the continuity of the callback path. A server makes a preliminary assessment of callback availability to a given client and avoids delegating responsibilities until it has determined that callbacks are supported. Because the granting of a delegation is always conditional upon the absence of conflicting access, clients must not assume that a delegation will be granted and they must always be prepared for OPENs to be processed without any delegations being granted.
Once granted, a delegation behaves in most ways like a lock. There is an associated lease that is subject to renewal together with all of the other leases held by that client.
Unlike locks, an operation by a second client to a delegated file will cause the server to recall a delegation through a callback.
On recall, the client holding the delegation must flush modified state (such as modified data) to the server and return the delegation. The conflicting request will not receive a response until the recall is complete. The recall is considered complete when the client returns the delegation or the server times out on the recall and revokes the delegation as a result of the timeout. Following the resolution of the recall, the server has the information necessary to grant or deny the second client's request.
At the time the client receives a delegation recall, it may have substantial state that needs to be flushed to the server. Therefore, the server should allow sufficient time for the delegation to be returned since it may involve numerous RPCs to the server. If the server is able to determine that the client is diligently flushing state to the server as a result of the recall, the server may extend the usual time allowed for a recall. However, the time allowed for recall completion should not be unbounded.
An example of this is when responsibility to mediate opens on a given file is delegated to a client (see the section "Open Delegation"). The server will not know what opens are in effect on the client. Without this knowledge the server will be unable to determine if the access and deny state for the file allows any particular open until the delegation for the file has been returned.
A client failure or a network partition can result in failure to respond to a recall callback. In this case, the server will revoke the delegation which in turn will render useless any modified state still on the client.
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There are three situations that delegation recovery must deal with:
In the event the client reboots or restarts, the failure to renew leases will result in the revocation of record locks and share reservations. Delegations, however, may be treated a bit differently.
There will be situations in which delegations will need to be reestablished after a client reboots or restarts. The reason for this is the client may have file data stored locally and this data was associated with the previously held delegations. The client will need to reestablish the appropriate file state on the server.
To allow for this type of client recovery, the server MAY extend the period for delegation recovery beyond the typical lease expiration period. This implies that requests from other clients that conflict with these delegations will need to wait. Because the normal recall process may require significant time for the client to flush changed state to the server, other clients need be prepared for delays that occur because of a conflicting delegation. This longer interval would increase the window for clients to reboot and consult stable storage so that the delegations can be reclaimed. For open delegations, such delegations are reclaimed using OPEN with a claim type of CLAIM_DELEGATE_PREV. (See the sections on "Data Caching and Revocation" and "Operation 18: OPEN" for discussion of open delegation and the details of OPEN respectively).
A server MAY support a claim type of CLAIM_DELEGATE_PREV, but if it does, it MUST NOT remove delegations upon SETCLIENTID_CONFIRM, and instead MUST, for a period of time no less than that of the value of the lease_time attribute, maintain the client's delegations to allow time for the client to issue CLAIM_DELEGATE_PREV requests. The server that supports CLAIM_DELEGATE_PREV MUST support the DELEGPURGE operation.
When the server reboots or restarts, delegations are reclaimed (using the OPEN operation with CLAIM_PREVIOUS) in a similar fashion to record locks and share reservations. However, there is a slight semantic difference. In the normal case if the server decides that a delegation should not be granted, it performs the requested action (e.g. OPEN) without granting any delegation. For reclaim, the server grants the delegation but a special designation is applied so that the client treats the delegation as having been granted but recalled by the server. Because of this, the client has the duty to write all modified state to the server and then return the delegation. This process of handling delegation reclaim reconciles three principles of the NFS version 4 protocol:
When a network partition occurs, delegations are subject to freeing by the server when the lease renewal period expires. This is similar to the behavior for locks and share reservations. For delegations, however, the server may extend the period in which conflicting requests are held off. Eventually the occurrence of a conflicting request from another client will cause revocation of the delegation. A loss of the callback path (e.g. by later network configuration change) will have the same effect. A recall request will fail and revocation of the delegation will result.
A client normally finds out about revocation of a delegation when it uses a stateid associated with a delegation and receives the error NFS4ERR_EXPIRED. It also may find out about delegation revocation after a client reboot when it attempts to reclaim a delegation and receives that same error. Note that in the case of a revoked write open delegation, there are issues because data may have been modified by the client whose delegation is revoked and separately by other clients. See the section "Revocation Recovery for Write Open Delegation" for a discussion of such issues. Note also that when delegations are revoked, information about the revoked delegation will be written by the server to stable storage (as described in the section "Crash Recovery"). This is done to deal with the case in which a server reboots after revoking a delegation but before the client holding the revoked delegation is notified about the revocation.
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When applications share access to a set of files, they need to be implemented so as to take account of the possibility of conflicting access by another application. This is true whether the applications in question execute on different clients or reside on the same client.
Share reservations and record locks are the facilities the NFS version 4 protocol provides to allow applications to coordinate access by providing mutual exclusion facilities. The NFS version 4 protocol's data caching must be implemented such that it does not invalidate the assumptions that those using these facilities depend upon.
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In order to avoid invalidating the sharing assumptions that applications rely on, NFS version 4 clients should not provide cached data to applications or modify it on behalf of an application when it would not be valid to obtain or modify that same data via a READ or WRITE operation.
Furthermore, in the absence of open delegation (see the section "Open Delegation") two additional rules apply. Note that these rules are obeyed in practice by many NFS version 2 and version 3 clients.
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For those applications that choose to use file locking instead of share reservations to exclude inconsistent file access, there is an analogous set of constraints that apply to client side data caching. These rules are effective only if the file locking is used in a way that matches in an equivalent way the actual READ and WRITE operations executed. This is as opposed to file locking that is based on pure convention. For example, it is possible to manipulate a two-megabyte file by dividing the file into two one-megabyte regions and protecting access to the two regions by file locks on bytes zero and one. A lock for write on byte zero of the file would represent the right to do READ and WRITE operations on the first region. A lock for write on byte one of the file would represent the right to do READ and WRITE operations on the second region. As long as all applications manipulating the file obey this convention, they will work on a local file system. However, they may not work with the NFS version 4 protocol unless clients refrain from data caching.
The rules for data caching in the file locking environment are:
Note that flushing data to the server and the invalidation of cached data must reflect the actual byte ranges locked or unlocked. Rounding these up or down to reflect client cache block boundaries will cause problems if not carefully done. For example, writing a modified block when only half of that block is within an area being unlocked may cause invalid modification to the region outside the unlocked area. This, in turn, may be part of a region locked by another client. Clients can avoid this situation by synchronously performing portions of write operations that overlap that portion (initial or final) that is not a full block. Similarly, invalidating a locked area which is not an integral number of full buffer blocks would require the client to read one or two partial blocks from the server if the revalidation procedure shows that the data which the client possesses may not be valid.
The data that is written to the server as a prerequisite to the unlocking of a region must be written, at the server, to stable storage. The client may accomplish this either with synchronous writes or by following asynchronous writes with a COMMIT operation. This is required because retransmission of the modified data after a server reboot might conflict with a lock held by another client.
A client implementation may choose to accommodate applications which use record locking in non-standard ways (e.g. using a record lock as a global semaphore) by flushing to the server more data upon an LOCKU than is covered by the locked range. This may include modified data within files other than the one for which the unlocks are being done. In such cases, the client must not interfere with applications whose READs and WRITEs are being done only within the bounds of record locks which the application holds. For example, an application locks a single byte of a file and proceeds to write that single byte. A client that chose to handle a LOCKU by flushing all modified data to the server could validly write that single byte in response to an unrelated unlock. However, it would not be valid to write the entire block in which that single written byte was located since it includes an area that is not locked and might be locked by another client. Client implementations can avoid this problem by dividing files with modified data into those for which all modifications are done to areas covered by an appropriate record lock and those for which there are modifications not covered by a record lock. Any writes done for the former class of files must not include areas not locked and thus not modified on the client.
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Client side data caching needs to respect mandatory file locking when it is in effect. The presence of mandatory file locking for a given file is indicated when the client gets back NFS4ERR_LOCKED from a READ or WRITE on a file it has an appropriate share reservation for. When mandatory locking is in effect for a file, the client must check for an appropriate file lock for data being read or written. If a lock exists for the range being read or written, the client may satisfy the request using the client's validated cache. If an appropriate file lock is not held for the range of the read or write, the read or write request must not be satisfied by the client's cache and the request must be sent to the server for processing. When a read or write request partially overlaps a locked region, the request should be subdivided into multiple pieces with each region (locked or not) treated appropriately.
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When clients cache data, the file data needs to be organized according to the file system object to which the data belongs. For NFS version 3 clients, the typical practice has been to assume for the purpose of caching that distinct filehandles represent distinct file system objects. The client then has the choice to organize and maintain the data cache on this basis.
In the NFS version 4 protocol, there is now the possibility to have significant deviations from a "one filehandle per object" model because a filehandle may be constructed on the basis of the object's pathname. Therefore, clients need a reliable method to determine if two filehandles designate the same file system object. If clients were simply to assume that all distinct filehandles denote distinct objects and proceed to do data caching on this basis, caching inconsistencies would arise between the distinct client side objects which mapped to the same server side object.
By providing a method to differentiate filehandles, the NFS version 4 protocol alleviates a potential functional regression in comparison with the NFS version 3 protocol. Without this method, caching inconsistencies within the same client could occur and this has not been present in previous versions of the NFS protocol. Note that it is possible to have such inconsistencies with applications executing on multiple clients but that is not the issue being addressed here.
For the purposes of data caching, the following steps allow an NFS version 4 client to determine whether two distinct filehandles denote the same server side object:
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When a file is being OPENed, the server may delegate further handling of opens and closes for that file to the opening client. Any such delegation is recallable, since the circumstances that allowed for the delegation are subject to change. In particular, the server may receive a conflicting OPEN from another client, the server must recall the delegation before deciding whether the OPEN from the other client may be granted. Making a delegation is up to the server and clients should not assume that any particular OPEN either will or will not result in an open delegation. The following is a typical set of conditions that servers might use in deciding whether OPEN should be delegated:
There are two types of open delegations, read and write. A read open delegation allows a client to handle, on its own, requests to open a file for reading that do not deny read access to others. Multiple read open delegations may be outstanding simultaneously and do not conflict. A write open delegation allows the client to handle, on its own, all opens. Only one write open delegation may exist for a given file at a given time and it is inconsistent with any read open delegations.
When a client has a read open delegation, it may not make any changes to the contents or attributes of the file but it is assured that no other client may do so. When a client has a write open delegation, it may modify the file data since no other client will be accessing the file's data. The client holding a write delegation may only affect file attributes which are intimately connected with the file data: size, time_modify, change.
When a client has an open delegation, it does not send OPENs or CLOSEs to the server but updates the appropriate status internally. For a read open delegation, opens that cannot be handled locally (opens for write or that deny read access) must be sent to the server.
When an open delegation is made, the response to the OPEN contains an open delegation structure which specifies the following:
The delegation stateid is separate and distinct from the stateid for the OPEN proper. The standard stateid, unlike the delegation stateid, is associated with a particular lock_owner and will continue to be valid after the delegation is recalled and the file remains open.
When a request internal to the client is made to open a file and open delegation is in effect, it will be accepted or rejected solely on the basis of the following conditions. Any requirement for other checks to be made by the delegate should result in open delegation being denied so that the checks can be made by the server itself.
The nfsace4 passed with delegation can be used to avoid frequent ACCESS calls. The permission check should be as follows:
The server may return an nfsace4 that is more restrictive than the actual ACL of the file. This includes an nfsace4 that specifies denial of all access. Note that some common practices such as mapping the traditional user "root" to the user "nobody" may make it incorrect to return the actual ACL of the file in the delegation response.
The use of delegation together with various other forms of caching creates the possibility that no server authentication will ever be performed for a given user since all of the user's requests might be satisfied locally. Where the client is depending on the server for authentication, the client should be sure authentication occurs for each user by use of the ACCESS operation. This should be the case even if an ACCESS operation would not be required otherwise. As mentioned before, the server may enforce frequent authentication by returning an nfsace4 denying all access with every open delegation.
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OPEN delegation allows much of the message overhead associated with the opening and closing files to be eliminated. An open when an open delegation is in effect does not require that a validation message be sent to the server. The continued endurance of the "read open delegation" provides a guarantee that no OPEN for write and thus no write has occurred. Similarly, when closing a file opened for write and if write open delegation is in effect, the data written does not have to be flushed to the server until the open delegation is recalled. The continued endurance of the open delegation provides a guarantee that no open and thus no read or write has been done by another client.
For the purposes of open delegation, READs and WRITEs done without an OPEN are treated as the functional equivalents of a corresponding type of OPEN. This refers to the READs and WRITEs that use the special stateids consisting of all zero bits or all one bits. Therefore, READs or WRITEs with a special stateid done by another client will force the server to recall a write open delegation. A WRITE with a special stateid done by another client will force a recall of read open delegations.
With delegations, a client is able to avoid writing data to the server when the CLOSE of a file is serviced. The file close system call is the usual point at which the client is notified of a lack of stable storage for the modified file data generated by the application. At the close, file data is written to the server and through normal accounting the server is able to determine if the available file system space for the data has been exceeded (i.e. server returns NFS4ERR_NOSPC or NFS4ERR_DQUOT). This accounting includes quotas. The introduction of delegations requires that a alternative method be in place for the same type of communication to occur between client and server.
In the delegation response, the server provides either the limit of the size of the file or the number of modified blocks and associated block size. The server must ensure that the client will be able to flush data to the server of a size equal to that provided in the original delegation. The server must make this assurance for all outstanding delegations. Therefore, the server must be careful in its management of available space for new or modified data taking into account available file system space and any applicable quotas. The server can recall delegations as a result of managing the available file system space. The client should abide by the server's state space limits for delegations. If the client exceeds the stated limits for the delegation, the server's behavior is undefined.
Based on server conditions, quotas or available file system space, the server may grant write open delegations with very restrictive space limitations. The limitations may be defined in a way that will always force modified data to be flushed to the server on close.
With respect to authentication, flushing modified data to the server after a CLOSE has occurred may be problematic. For example, the user of the application may have logged off the client and unexpired authentication credentials may not be present. In this case, the client may need to take special care to ensure that local unexpired credentials will in fact be available. This may be accomplished by tracking the expiration time of credentials and flushing data well in advance of their expiration or by making private copies of credentials to assure their availability when needed.
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When a client holds a write open delegation, lock operations are performed locally. This includes those required for mandatory file locking. This can be done since the delegation implies that there can be no conflicting locks. Similarly, all of the revalidations that would normally be associated with obtaining locks and the flushing of data associated with the releasing of locks need not be done.
When a client holds a read open delegation, lock operations are not performed locally. All lock operations, including those requesting non-exclusive locks, are sent to the server for resolution.
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The server needs to employ special handling for a GETATTR where the target is a file that has a write open delegation in effect. The reason for this is that the client holding the write delegation may have modified the data and the server needs to reflect this change to the second client that submitted the GETATTR. Therefore, the client holding the write delegation needs to be interrogated. The server will use the CB_GETATTR operation. The only attributes that the server can reliably query via CB_GETATTR are size and change.
Since CB_GETATTR is being used to satisfy another client's GETATTR request, the server only needs to know if the client holding the delegation has a modified version of the file. If the client's copy of the delegated file is not modified (data or size), the server can satisfy the second client's GETATTR request from the attributes stored locally at the server. If the file is modified, the server only needs to know about this modified state. If the server determines that the file is currently modified, it will respond to the second client's GETATTR as if the file had been modified locally at the server.
Since the form of the change attribute is determined by the server and is opaque to the client, the client and server need to agree on a method of communicating the modified state of the file. For the size attribute, the client will report its current view of the file size. For the change attribute, the handling is more involved.
For the client, the following steps will be taken when receiving a write delegation:
For simplicity of implementation, the client MAY for each CB_GETATTR return the same value d. This is true even if, between successive CB_GETATTR operations, the client again modifies in the file's data or metadata in its cache. The client can return the same value because the only requirement is that the client be able to indicate to the server that the client holds modified data. Therefore, the value of d may always be c + 1.
While the change attribute is opaque to the client in the sense that it has no idea what units of time, if any, the server is counting change with, it is not opaque in that the client has to treat it as an unsigned integer, and the server has to be able to see the results of the client's changes to that integer. Therefore, the server MUST encode the change attribute in network order when sending it to the client. The client MUST decode it from network order to its native order when receiving it and the client MUST encode it network order when sending it to the server. For this reason, change is defined as an unsigned integer rather than an opaque array of octets.
For the server, the following steps will be taken when providing a write delegation:
As discussed earlier in this section, the client MAY return the same cc value on subsequent CB_GETATTR calls, even if the file was modified in the client's cache yet again between successive CB_GETATTR calls. Therefore, the server must assume that the file has been modified yet again, and MUST take care to ensure that the new nsc it constructs and returns is greater than the previous nsc it returned. An example implementation's delegation record would satisfy this mandate by including a boolean field (let us call it "modified") that is set to false when the delegation is granted, and an sc value set at the time of grant to the change attribute value. The modified field would be set to true the first time cc != sc, and would stay true until the delegation is returned or revoked. The processing for constructing nsc, time_modify, and time_metadata would use this pseudo code:
if (!modified) {
do CB_GETATTR for change and size;
if (cc != sc)
modified = TRUE;
} else {
do CB_GETATTR for size;
}
if (modified) {
sc = sc + 1;
time_modify = time_metadata = current_time;
update sc, time_modify, time_metadata into file's metadata;
}
return to client (that sent GETATTR) the attributes
it requested, but make sure size comes from what
CB_GETATTR returned. Do not update the file's metadata
with the client's modified size.
In the case that the file attribute size is different than the server's current value, the server treats this as a modification regardless of the value of the change attribute retrieved via CB_GETATTR and responds to the second client as in the last step.
This methodology resolves issues of clock differences between client and server and other scenarios where the use of CB_GETATTR break down.
It should be noted that the server is under no obligation to use CB_GETATTR and therefore the server MAY simply recall the delegation to avoid its use.
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The following events necessitate recall of an open delegation:
Whether a RENAME of a directory in the path leading to the file results in recall of an open delegation depends on the semantics of the server file system. If that file system denies such RENAMEs when a file is open, the recall must be performed to determine whether the file in question is, in fact, open.
In addition to the situations above, the server may choose to recall open delegations at any time if resource constraints make it advisable to do so. Clients should always be prepared for the possibility of recall.
When a client receives a recall for an open delegation, it needs to update state on the server before returning the delegation. These same updates must be done whenever a client chooses to return a delegation voluntarily. The following items of state need to be dealt with:
In the case of write open delegation, file locking imposes some additional requirements. To precisely maintain the associated invariant, it is required to flush any modified data in any region for which a write lock was released while the write delegation was in effect. However, because the write open delegation implies no other locking by other clients, a simpler implementation is to flush all modified data for the file (as described just above) if any write lock has been released while the write open delegation was in effect.
An implementation need not wait until delegation recall (or deciding to voluntarily return a delegation) to perform any of the above actions, if implementation considerations (e.g. resource availability constraints) make that desirable. Generally, however, the fact that the actual open state of the file may continue to change makes it not worthwhile to send information about opens and closes to the server, except as part of delegation return. Only in the case of closing the open that resulted in obtaining the delegation would clients be likely to do this early, since, in that case, the close once done will not be undone. Regardless of the client's choices on scheduling these actions, all must be performed before the delegation is returned, including (when applicable) the close that corresponds to the open that resulted in the delegation. These actions can be performed either in previous requests or in previous operations in the same COMPOUND request.
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A client may fail to respond to a recall for various reasons, such as a failure of the callback path from server to the client. The client may be unaware of a failure in the callback path. This lack of awareness could result in the client finding out long after the failure that its delegation has been revoked, and another client has modified the data for which the client had a delegation. This is especially a problem for the client that held a write delegation.
The server also has a dilemma in that the client that fails to respond to the recall might also be sending other NFS requests, including those that renew the lease before the lease expires. Without returning an error for those lease renewing operations, the server leads the client to believe that the delegation it has is in force.
This difficulty is solved by the following rules:
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At the point a delegation is revoked, if there are associated opens on the client, the applications holding these opens need to be notified. This notification usually occurs by returning errors for READ/WRITE operations or when a close is attempted for the open file.
If no opens exist for the file at the point the delegation is revoked, then notification of the revocation is unnecessary. However, if there is modified data present at the client for the file, the user of the application should be notified. Unfortunately, it may not be possible to notify the user since active applications may not be present at the client. See the section "Revocation Recovery for Write Open Delegation" for additional details.
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When locks and delegations are revoked, the assumptions upon which successful caching depend are no longer guaranteed. For any locks or share reservations that have been revoked, the corresponding owner needs to be notified. This notification includes applications with a file open that has a corresponding delegation which has been revoked. Cached data associated with the revocation must be removed from the client. In the case of modified data existing in the client's cache, that data must be removed from the client without it being written to the server. As mentioned, the assumptions made by the client are no longer valid at the point when a lock or delegation has been revoked. For example, another client may have been granted a conflicting lock after the revocation of the lock at the first client. Therefore, the data within the lock range may have been modified by the other client. Obviously, the first client is unable to guarantee to the application what has occurred to the file in the case of revocation.
Notification to a lock owner will in many cases consist of simply returning an error on the next and all subsequent READs/WRITEs to the open file or on the close. Where the methods available to a client make such notification impossible because errors for certain operations may not be returned, more drastic action such as signals or process termination may be appropriate. The justification for this is that an invariant for which an application depends on may be violated. Depending on how errors are typically treated for the client operating environment, further levels of notification including logging, console messages, and GUI pop-ups may be appropriate.
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Revocation recovery for a write open delegation poses the special issue of modified data in the client cache while the file is not open. In this situation, any client which does not flush modified data to the server on each close must ensure that the user receives appropriate notification of the failure as a result of the revocation. Since such situations may require human action to correct problems, notification schemes in which the appropriate user or administrator is notified may be necessary. Logging and console messages are typical examples.
If there is modified data on the client, it must not be flushed normally to the server. A client may attempt to provide a copy of the file data as modified during the delegation under a different name in the file system name space to ease recovery. Note that when the client can determine that the file has not been modified by any other client, or when the client has a complete cached copy of file in question, such a saved copy of the client's view of the file may be of particular value for recovery. In other case, recovery using a copy of the file based partially on the client's cached data and partially on the server copy as modified by other clients, will be anything but straightforward, so clients may avoid saving file contents in these situations or mark the results specially to warn users of possible problems.
Saving of such modified data in delegation revocation situations may be limited to files of a certain size or might be used only when sufficient disk space is available within the target file system. Such saving may also be restricted to situations when the client has sufficient buffering resources to keep the cached copy available until it is properly stored to the target file system.
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The attributes discussed in this section do not include named attributes. Individual named attributes are analogous to files and caching of the data for these needs to be handled just as data caching is for ordinary files. Similarly, LOOKUP results from an OPENATTR directory are to be cached on the same basis as any other pathnames and similarly for directory contents.
Clients may cache file attributes obtained from the server and use them to avoid subsequent GETATTR requests. Such caching is write through in that modification to file attributes is always done by means of requests to the server and should not be done locally and cached. The exception to this are modifications to attributes that are intimately connected with data caching. Therefore, extending a file by writing data to the local data cache is reflected immediately in the size as seen on the client without this change being immediately reflected on the server. Normally such changes are not propagated directly to the server but when the modified data is flushed to the server, analogous attribu