Transaction log management in a disconnectable computer and network6192365Abstract A method and apparatus are disclosed for managing a transaction log which contains updates representing operations performed on a database replica in a network of disconnectable computers. The invention provides for compression of the log by the identification and removal of redundant updates. Log compression removes apparent inconsistencies between operations performed on disconnected computers, reduces storage requirements on each computer, and speeds up transaction synchronization when the computers are reconnected. The invention also provides for restoration of prior versions of database objects using the log. Claims What is claimed and desired to be secured by patent is: Description FIELD OF THE INVENTION
CLASS ha_server
{
SUPERCLASS ndr_dodb_object_header;
PARENT ndr_dodb_database;
PROPERTY NDR_OS_CLASS_FLAG_FULLY_REPLICATED;
ATTRIBUTE
{
ha_server_name server_name
PROPERTY NDR_OS_ATTR_FLAG_SIBLING_KEY;
}
}
CONSTANT HA_VOLUME_NAME_MAX = 32;
DATATYPE ha_volume_name STRING_HA_VOLUME_NAME_MAX;
DATATYPE ha_volume_id BYTE;
A volume has a name, which must be unique within the server and can be used as the root component of a path name:
CLASS ha_volume
{
SUPERCLASS ndr_dodb_object_header;
PARENT ha_server;
PROPERTY NDR_OS_CLASS_FLAG_NAMESPACE_ROOT;
ATTRIBUTE
{
ha_volume_name volume_name
PROPERTY NDR_OS_ATTR_FLAG_SIBLING_KEY .linevert
split.
NDR_OS_ATTR_FLAG_IS_DOS_FILENAME;
ha_volume_id volume_id;
}
}
In order to allocate unique volume identifiers this object holds the next free volume ID. Initially this is set to 1, so that the SYS volume can be given ID 0 when it is added to the database, in case any applications make assumptions about SYS:
CLASS ha_next_volume
{
SUPERCLASS ndr_dodb_object_header;
PARENT ha_server;
PROPERTY NDR_OS_CLASS_FLAG_UNREPLICATED;
ATTRIBUTE
{
ndr_dodb_dummy_key dummy_key
PROPERTY NDR_OS_ATTR_FLAG_SIBLING_KEY
COMPARISON ndr_dodb_dummy_key_compare
VALIDATION ndr_dodb_dummy_key_validate;
ha_volume_id next_free_volume_id;
}
}
A file or directory name can be 12 (2-byte) characters long: CONSTANT HA_FILENAME_MAX=24; DATATYPE ha_filename STRING HA_FILENAME_MAX; The ha_file_or_dir_id is a compound unique key embracing the file or directory ID that is allocated by the server, as well as the server-generated volume number. The latter is passed as a byte from class 87 NetWare Core Protocols from which it is read directly into vol (declared as a byte below). Elsewhere in the code the type ndr_host_volume_id (a UINT16) is used for the same value.
DATATYPE ha_file_or_dir_id
ULONG file_or_dir;
ha_volume_id vol;
}
Files and directories have many shared attributes, the most important being the file name. This must be unique for any parent directory.
CLASS ha_file_or_dir
{
PARENT ha_directory;
SUPERCLASS ndr_dodb_object_header;
ATTRIBUTE
{
ha_filename filename
PROPERTY NDR_OS_ATTR_FLAG_SIBLING_KEY .linevert split.
NDR_OS_ATTR_FLAG_IS_DOS_FILENAME;
ha_file_or_dir_id id
PROPERTY NDR_OS_ATTR_FLAG_GLOBAL_KEY .linevert split.
NDR_OS_ATTR_FLAG_UNREPLICATED
GROUP file_or_dir_id_group;
ULONG attributes;
SHORT creation_date;
SHORT creation_time;
ndr_dodb_auth_id creation_id;
SHORT access_date;
SHORT archive_date;
SHORT archive_time;
ndr_dodb_auth_id archive_id;
}
}
A file has some additional attributes not present in a directory, and may contain a contents fork which can be accessed via a file distributor 90 (FIG. 3):
CLASS ha_file
{
SUPERCLASS ha_file_or_dir;
PROPERTY NDR_OS_CLASS_FLAG_DEFINE_REPLICAS .linevert split.
NDR_OS_CLASS_FLAG_HAS_PARTIALLY_REPLICATED_FILE
NDR_OS_CLASS_FLAG_HAS_FILE_PATH_NAME .linevert split.
NDR_OS_CLASS_FLAG_PARENT_HAS_RSC;
ATTRIBUTE
{
BYTE execute_type;
SHORT update_date
property NDR_OS_ATTR_FLAG_UNREPLICATED;
SHORT update_time
property NDR_OS_ATTR_FLAG_UNREPLICATED;
ndr_dodb_auth_id update_id
property NDR_OS_ATTR_FLAG_UNREPLICATED;
ULONG length
property NDR_OS_ATTR_FLAG_UNREPLICATED;
}
}
A directory does not possess a contents fork for file distributor 90 access. The access rights mask is inherited and should be managed by like access control lists ("ACLs"):
CLASS ha_directory
{
SUPERCLASS ha_file_or_dir;
PROPERTY NDR_OS_CLASS_FLAG_DEFINE_REPLICAS .linevert split.
NDR_OS_CLASS_FLAG_HAS_FILE_PATH_NAME .linevert split.
NDR_OS_CLASS_FLAG_HAS_RSC;
//replication support count
ATTRIBUTE
{
BYTE access_rights_mask;
SHORT update_date;
SHORT update_time;
ndr_dodb_auth_id update_id;
SHORT rsc
PROPERTY NDR_OS_ATTR_FLAG IS_RSC .linevert split.
NDR_OS_ATTR_FLAG_UNREPLICATED;
}
}
The root directory must appear at the top of the hierarchy below the volume. Its name is not used; the volume name is used instead. This is the top of the replication hierarchy and therefore is the top level RSC in this hierarchy:
CLASS ha_root_directory
{
SUPERCLASS ha_directory;
PARENT ha_volume;
PROPERTY NDR_OS_CLASS_FLAG_DEFINE.sub.--
REPLICAS .linevert split.
NDR_OS_CLASS_FLAG_HAS_RSC;
}
In one embodiment, schemas such as the schema 84 are defined in a source code format and then compiled to generate C language header files and tables. The named source file is read as a stream of lexical tokens and parsed using a recursive descent parser for a simple LL(1) syntax. Parsing an INCLUDE statement causes the included file to be read at that point. Once a full parse tree has been built (using binary nodes), the tree is walked to check for naming completeness. The tree is next walked in three passes to generate C header (.H) files for each included schema file. The header generation passes also compute information (sizes, offsets, and so forth) about the schema which is stored in Id nodes in the tree. Finally, the complete tree is walked in multiple passes to generate the schema table C source file, which is then ready for compiling and linking into an agent's executable program. Each disconnectable computer 40 also includes a replica manager 46 which initiates and tracks location-specific updates as necessary in response to database manager 42 requests. The replica manager is discussed in detail in connection with later Figures. A file system interface 48 on each computer 40 mediates between the replica manager 46 and a storage device and controller 54. Suitable file system interfaces 48 include well-known interfaces 48 such as the File Allocation Table ("FAT") interfaces of various versions of the MS-DOS.RTM. operating system (MS-DOS is a registered trademark of Microsoft Corporation), the XENIX.RTM. file system (registered trademark of Microsoft Corporation), the various NOVELL file systems (trademark of Novell, Inc.), the various UNIX file systems (trademark of Santa Cruz Operations), the PCIX file system, the High Performance File System ("HPFS") used by the OS/2 operating system (OS/2 is a mark of International Business Machines Corporation), and other conventional file systems. Suitable storage devices and respective controllers 54 include devices and controllers for the media disclosed above in connection with the storage medium 34 (FIG. 1) and other conventional devices and controllers, including non-volatile storage devices. It is understood, however, that the database replicas 56 stored on these media are not necessarily conventional even though the associated devices and controllers 54 may themselves be known in the art. Each computer 40 also has a network link manager 50 that is capable of establishing a network connection 52 with another disconnectable computer 40. Suitable network link managers 50 include those capable of providing remote procedure calls or an equivalent communications and control capability. One embodiment utilizes "DataTalk" remote procedure call software with extended NetWare Core Protocol calls and provides functionality according to the following interface:
rpc_init() Initialize RPC subsystem
rpc_shutdown() Shutdown RPC subsystem
rpc_execute() Execute request at single
location
rpc_ping() Ping a location (testing)
rpc_claim_next_execute() Wait until the next rpc_execute()
is guaranteed to be used by this
thread
rpc_free_next_execute() Allow others to use rpc_execute()
Those of skill in the art will appreciate that other remote procedure call mechanisms may also be employed according to the present invention. Suitable network connections 52 may be established using packet-based, serial, internet-compatible, local area, metropolitan area, wide area, and wireless network transmission systems and methods. FIGS. 2 and 3 illustrate one embodiment of the replica manager 46 of the present invention. A replica distributor 70 insulates the database manager 42 from the complexities caused by having database entries stored in replicas 56 on multiple computers 40 while still allowing the database manager 42 to efficiently access and manipulate individual database objects, variables, and/or records. A replica processor 72 maintains information about the location and status of each replica 56 and ensures that the replicas 56 tend to converge. A consistency distributor 74 and a consistency processor 76 cooperate to maintain convergent and transactional consistency of the database replicas 56. The major processes used include an update process which determines how transaction updates are applied, an asynchronous synchronization process that asynchronously synchronizes other locations in a location set, a synchronous synchronization process that synchronously forces two locations into sync with each other, an optional concurrency process that controls distributed locking, and a merge process that adds new locations to a location set. In one embodiment, processes for synchronization and merging are implemented using background software processes with threads or similar means. The concurrency process may be replaced by a combination of retries and clash handling to reduce implementation cost and complexity. Each location is identified by a unique location identifier. A "location sync group" is the group of all locations that a specific location synchronizes with. The location sync group for a database replica 56 on a client 20 is the client and the server 16 or other computer 28 that holds a master replica 56; the computer 28 holding the master replica 56 is the "storage location" of the target database. The location sync group for the computer 28 that holds the master replica 56 is all computers 28 connectable to the network that hold a replica 56. A "location set" is a set of presently connected locations in a location sync group. Locations in an "active location set" have substantially converged, while those in a "merge location set" are currently being merged into the active location set. Objects are read at a "reference location" and updated at an "update location," both of which are local when possible for performance reasons. To support concurrency control, objects require a "lock location" where they are locked for read or update; the local location is the same for all processes in a given location set. According to one update process of the present invention, the updates for a single transaction are all executed at one update location. Each group of updates associated with a single transaction have a processor transaction identifier ("PTID") containing the location identifier of the update location and a transaction sequence number. The transaction sequence number is preferably monotonically consecutively increasing for all completed transactions at a given location, even across computer 28 restarts, so that other locations receiving updates can detect missed updates. The PTID is included in update details written to an update log by an object processor 86. An update log (sometimes called an "update stream") is a chronological record of operations on the database replica 56. Although it may be prudent to keep a copy of an update log on a non-volatile storage device, this is not required. The operations will vary according to the nature of the database, but typical operations include adding objects, removing objects, modifying the values associated with an object attribute, modifying the attributes associated with an object, and moving objects relative to one another. The PTID is also included as an attribute of each target database object to reflect the latest modification of the object. In one embodiment, the PTID is also used to create a unique (within the target database) unique object identifier ("UOID") when a target database object is first created. A target database object may contain attributes that can be independently updated. For instance, one user may set an archive attribute on a file while a second user independently renames the file. In such situations, an object schema 84 may define attribute groups. A separate PTID is maintained in the object for each attribute group, thereby allowing independent updates affecting different attribute groups of an object to be automatically merged without the updates being treated as a clash. The consistency distributor 74 gathers all of the updates for a single transaction and sends them, at close transaction time, to the update location for the transaction. The consistency processor 76 on the update location writes the updates to a transaction logger 88. In one embodiment, the transaction logger 88 buffers the updates in memory (e.g. RAM). If the update location is not local then the updates are committed to the transaction log and the PTID for the transaction is returned, so that the same updates can be buffered locally; this allows all updates to be applied in order locally. In this manner the transaction updates are applied to the update location. An objective of one asynchronous synchronization process of the present invention is to keep the rest of the locations in the location set in sync without unacceptable impact on foreground software process performance. This is achieved by minimizing network transfers. A process of the consistency processor 76 (such as a background software process) either periodically or on demand requests the transaction logger 88 to force write all pending transactions to the log and (eventually) to the target database. The consistency processor 76 also causes the batch of updates executed at an update location to be transmitted to all other locations in the current location set as a "SyncUpdate" request. These updates are force written to the log before they are transmitted to other locations, thereby avoiding use of the same transaction sequence number for different transactions in the event of a crash. The SyncUpdate requests are received by other locations in the same location set and applied to their in-memory transaction logs by their respective consistency processors 76. Each consistency processor 76 only applies SyncUpdate transactions which have sequence numbers that correspond to the next sequence number for the specified location. The consistency processor 76 can determine if it has missed updates or received them out of order by examining the PTID. If updates are missed, the PTID of the last transaction properly received is sent to the consistency distributor 74 that sent out the updates, which then arranges to send the missing updates to whichever consistency processors 76 need them. Acknowledged requests using threads or a similar mechanism can be used in place of unacknowledged requests sent by non-central locations. Non-central locations (those not holding a master replica 56) only need to synchronize with one location and thus only require a small number of threads. To promote scalability, however, central locations preferably use unacknowledged broadcasts to efficiently transmit their SyncUpdate requests. The asynchronous synchronization process causes SyncUpdate requests to be batched to minimize network transfers. However, the cost paid is timeliness. Accordingly, a synchronous synchronization process according to the present invention may be utilized to selectively speed up synchronization. The synchronous synchronization process provides a SyncUptoPTID request and response mechanism. In one embodiment, the SyncUptoPTID mechanism utilizes a SyncState structure which is maintained as part of a location state structure or location list that is managed by a location state processor 80 in the memory of each computer 28. The SyncState structure for a given location contains a location identifier and corresponding transaction sequence number for the most recent successful transaction applied from that location. The SyncState structure is initialized from the update log at startup time and updated in memory as new transactions are applied. A SyncUptoPTID request asks a destination to bring itself up to date with a source location according to a PTID. The destination sends a copy of the SyncState structure for the source location to that source location. The source location then sends SyncUpdate requests to the destination location, as previously described, up to an including the request with the PTID that was specified in the SyncUptoPTID request. In a preferred embodiment, the central server is a NetWare server and the SyncUptoPTID requirements are approximately 100 bytes per location, so scalability is not a significant problem for most systems. A merge process according to the present invention includes merging location sets when disconnected disconnectable computers are first connected or reconnected. For instance, merging location sets normally occurs when a computer new to the network starts up and merges into an existing location set. Merging can also happen when two sets of computers become connected, such as when a router starts. Merging occurs when two replicas 56 are resynchronized after the computers 28 on which the replicas 56 reside are reconnected following a period of disconnection. Either or both of the computers 28 may have been shut down during the disconnection. A set of updates are "merged atomically" if they are merged transactionally on an all-or-nothing basis. A distributed database is "centrally synchronized" if one computer 28, sometimes denoted the "central server," carries a "master replica" with which all merges are performed. Portions of the master replica or portions of another replica 56 may be "shadowed" during a merge. A shadow replica, sometimes called a "shadow database", is a temporary copy of at least a portion of the database. The shadow database is used as a workspace until it can be determined whether changes made in the workspace are consistent and thus can all be made in the shadowed replica, or are inconsistent and so must all be discarded. The shadow database uses an "orthogonal name space." That is, names used in the shadow database follow a naming convention which guarantees that they will never be confused with names in the shadowed database. A "state-based" approach to merging compares the final state of two replicas 56 and modifies one or both replicas 56 to make corresponding values equal. A "log-based" or "transaction-based" approach to merging incrementally applies successive updates made on a first computer 28 to the replica 56 stored on a second computer 28, and then repeats the process with the first computer's replica 56 and the second computer's update log. A hybrid approach uses state comparison to generate an update stream that is then applied incrementally. The present invention preferably utilizes transaction-based merging rather than state-based or hybrid merging. As an illustration, consider the process of merging a single new location A with a location set containing locations B and C. In one embodiment, the following performance goals are satisfied: (a) Use of locations B and C is not substantially interrupted by synchronization of the out-of-date location A with B and C; and (b) Users connected to location A (possibly including multiple users if location B is a gateway) are able to see the contents of the other locations in the set within a reasonable period of time. Merging typically occurs in three phases. During a "merging out" phase location A sends newer updates to location B. For instance, if A's location list contains PTID 50:14 (location identifier:transaction sequence number) and B's location list contains PTID 50:10, then the newer updates sent would correspond to PTID values 50:11 through 50:14. During a "merging in" phase new updates in the merge location B are merged into A's location. For instance, suppose A's location list contains PTIDs 100:12 and 150:13 and B's location list contains PTIDs 100:18 and 150:13. Then the new updates would correspond to PTID values 100:13 through 100:18. If updates are in progress when merging is attempted, the initial attempt to merge will not fully succeed, and additional iterations of the merging in and merging out steps are performed. In one embodiment, merging does not include file contents synchronization. Instead file contents are merged later, either by a background process or on demand triggered by file access. This reduces the time required for merging and promotes satisfaction of the two performance goals identified above. In embodiments tailored to "slow" links, merging is preferably on-going to take advantage of whatever bandwidth is available without substantially degrading the perceived performance of other processes running on the disconnectable computers. In embodiments employing an update log, the log is preferably compressed prior to merging. Compression reduces the number of operations stored in the log. Compression may involve removing updates from the log, altering the parameters associated with an operation in a given update, and/or changing the order in which updates are stored in the log. In one embodiment, all Object Database calls come through the consistency distributor 74, which manages distributed transaction processing and maintains consistency between locations. Almost all calls from a location distributor 78 are made via the consistency distributor 74 because the consistency distributor 74 supports a consistent view of the locations and the database replicas 56 on them. The consistency distributor 74 and an object distributor 82 support multiple concurrent transactions. This is needed internally to allow background threads to be concurrently executing synchronization updates. It could also be used to support multiple concurrent gateway users. In an alternative embodiment, multiple concurrent transactions on the same session is supported through the consistency distributor 74. In one embodiment, the consistency distributor 74 and the consistency processor 76 are implemented in the C programming language as a set of files which provide the functionality described here. Files CD.H and CD.C implement part of the consistency distributor 74. A separate module having files CD_BG.H and CD_BG.C is responsible for background processes associated with merging and synchronization. A module having files CDI.H and CDI.C contains functions used by both the CD and CD_BG modules. These modules provide functionality according to the following interface:
cd_init Init CD
cd_shutdown Shutdown CD
cd_create_replica Create a replica of a specified
database
cd_remove_replica Remove a replica of a
specified database
cd_load_db Load an existing database
cd_unload_db Unload an existing database
cd_merge_start Start merge of active and
merge location sets
cd_merge_stop Stop merge
cd_start_txn Start a CD transaction
cd_set_txn_ref_loc Set reference/update lid
(location identifier) for txn
(transaction)
cd_get_txn_desc Get a txn descriptor given
a txn id
cd_abort_txn Abort a CD transaction
cd_end_txn End a CD transaction
cd_commit Commit all previously closed
txns to disk
cd_execute_txn Execute locks and updates
for a txn
cd_read Do read or lookup request
cd_readn Do readn
cd_lookup_by_uoid Do lookup using UOID
cd_add_lock Add an object or agent lock
cd_remove_lock Remove an object or agent
lock
cd_modify_attribute Modify a single attribute in a
previously read object
cd_init_new_doid Setup all fields in a new doid
cd_add Add a new object
cd_remove Remove an object
cd_move Move an object
cd_set_marker Add marker point to txn
cd_revert_to_marker Revert txn state to last marker
cd_get_effective_access_right Get the effective access rights
for the current session and
object
cd_convert_uoid2doid Convert UOID to DOID
cd_sync_object Get the server to send a newly
replicated object
cd_bg_init Initialize CD background
processes
cd_bg_merge Execute a background
merge
cd_bg_sync_remote_upto_ptid Bring remote location up
to date with local PTID
cdi_init
cdi_shutdown
cdi_execute_ack_sys Execute acknowledged
request using system
session
cdi_execute_ack Execute acknowledged
request
cdi_apply_locks Apply locks for txn
cdi_abort_prc_txn Remove all locks already
set for a txn
//Forced update location (used to change update location
when executing clash handler functions)
cdi_register_forced_update_location Register location to be used as
update location for thread
cdi_unregister_forced_update_location Unregister location to be
used as update location
for thread
cdi_get_forced_update_location Get forced update location for
thread
cdi_sync_upto_ptid Bring location up to date
with PTID
cdi_sync_upto_now Bring location up to date
with latest PTID
cdi_sync_loc_list Make my location list
consistent with destination
location list and return
info on mismatch of PTIDs
cdi_read_loc_list Read location list
cdi_sync_upto_dtid Bring location up to date with
DTID
Since updates are cached during a transaction, special handling of reads performed when updates are cached is required. In one embodiment, the caller of cd_read() or cd_readn() sees the results of all updates previously executed in the transaction. In an alternative embodiment, for cd_read() reads will see all previously added objects and will see the modified attributes of objects, but will not see the effects of moves or removes. Thus if an object is removed during a transaction the read will behave as if it has not been removed. The same is true for moved objects. Modifications to keys will have no effect on reads using the keys. The cd_readn() function behaves as if none of the updates in the current transaction have been applied. In one embodiment, the consistency processor 76, which processes all distributed object database requests, includes background processes that manage object database updates on local locations and synchronization of locations. Within this embodiment, a CP module contains a dispatcher for all requests which call functions that have a prefix of "cpXX_"; a CPR module processes read requests; a CPU module processes update and lock requests; a CPSM module processes synchronization and merging requests; a CP_BG module controls background processing which includes scheduling multiple background threads, controlling the state of all local locations and synchronization of local locations with local and remote locations; and a CPUI module provides functions that are shared by the CP_BG and CPx modules. These modules provide functionality according to the following interface:
cp_init Includes performing mounting of
local locations and recovery of
TL (transaction logger 88) and OP
(object processor 86)
cp_shutdown Shutdown CP
cp_process Process a consistency request
cp_clear_stats Reset CP statistics
cp_dump_stats Dump CP statistics to the log
cpr_process_read Process OP read or lookup request
cpr_process_readn Process readn request
cpu_register_dtid Reqister use of a DTID at a
reference location
cpu_execute_txn Execute single txn at reference
location
cpu_commit Commit all txns for session
cpu_add_locks Add list of locks
cpu_remove_locks Remove list of locks
cpu_abort_prc_txn Remove object locks for specified
transaction
cpsm_sync_upto_ptid Bring remote locations up to date
as far as given PTID
cpsm_get_latest_ptid Obtain the latest PTID
cpsm_get_sync_object Remote machine wants to sync a
newly replicated object
cpsm_sync_object Add a newly replicated object to
the local database
cpsm_get_sync_update Get a local sync_update
cpsm_sync_update Apply multiple update txns to
location
cpsm_read_loc_list Read list of locations and states
cpsm_sync_loc_list Sync location list
cpsm_merge_loc_list Attempt to merge my location list
with other location list
cpsm_sync_finished Remote machine is notifying us
that a sync_upto_ptid has
completed
cpsm_request_merge Request a merge of this location
with the central server
cpui_init Initialize internal structures
cpui_shutdown Shutdown CPUI subsystem
cpui_execute_txn Execute update txn at a local
location
cpui_apply_update_list_to_db Apply an update list to an OP
database
cpui_commit Commit all txns at location
cpui_flush Flush all txns to object database
at location
cpui_replay_logged_transactions Replay transactions from the log
that have not been committed to
OP
cp_bg_init Initialize CP_BG subsystem
cp_bg_shutdown Shutdown CP_BG subsystem
cp_bg_handle_distributed_request Handle a request that requires
remote communication
cp_bg_notify_close_txn Notify CP_BG of a closed
transaction
cp_bg_notify_commit Notify CP_BG that all txns are
committed at a location
cp_bg_attempt_send_flush Attempt to send out and flush
txns
cp_bg_notify_load Notify CP_BG of a newly loaded
DB
cp_bg_notify_unload Notify CP_BG of a newly
unloaded DB
cp_bg_flush upto_ptid Force all transactions upto the
specified ptid to the migrated
state
The location distributor 78 in each replica manager 46 and the location state processor 80 are used to determine the storage locations of database entries. In one embodiment, the location state processor 80 uses a cache of the current state of locations and maintains state information on the merging process. The location state processor 80 is responsible for processing remote requests which pertain to the location list. All locations that are up at any time within a sync group are in either the ACTIVE or MERGE location sets. The ACTIVE location set contains all locations that are in sync with the local location up to certain sync watermarks. The MERGE location set contains all nodes that are not in sync with the local location, either through not having updates the active set does have, or through having updates the active set does not have. Locations in the MERGE set enter the ACTIVE set through the two-way merging process described above, under control of the consistency distributor 74 and the consistency processor 76. Once in the ACTIVE set, a location should never leave it until the location goes down. Each location continuously sends out its local updates to other members of its active location set as part of the merging process. The PTID in a location's log that was last sent out in this manner is called the location's "low watermark" PTID. For a location to enter the active set it must have all PTIDS in its local log up to the low watermark PTID; only the merging process used to move a location from the MERGE to the ACTIVE location set is capable of propagating early transactions. Each location also maintains a "high watermark" PTID which is the last transaction (in local log order) that has been committed, and is thus a candidate for sending out in a background sync update. The replica managers 46 track the last transaction sequence number made by every location up to the low watermark PTID in order to know whether a location is up to date with another location's low watermark. The log ordering may be different in different locations, up to an interleave. One embodiment of the location state processor 80 provides functionality according to the following interface:
ls_init Initialize LS
ls_shutdown Shutdown LS
ls_close_db Clear out all entries for a
database
ls_a11ocate_new_lid Allocate a new location
identifier for use by a new
replica
ls_add Add a new location
ls_remove Remove a location
ls_modify_local_tid Modify a location entry's local
transaction identifier (sequence
number)
ls_modify_state Modify a location entry's state
ls_get_loc_list Get list of locations
ls_get_loc_sync_list Get list of locations for syncing
ls_get_next_loc Get next location
ls_get_first_in_loc_list Get first location in list that
is in current location set
ls_get_loc_entry Get location entry given lid
(location identifier)
ls_get_first_ref_loc Get nearest reference location in
provided list
ls_get_first_ref_loc_in_list Get first reference location in
provided list
ls_get_lock_loc Get lock location for location
set
ls_higher_priority Determine which location has
highest priority
ls_complete_merge Complete the merge process
ls_set_sync_watermarks Set the high and low watermark
PTIDS used in syncing and merging
The object distributor 82 manages ACLs and otherwise manages access to objects in the database. In one embodiment, the object distributor 82 provides functionality according to this interface:
typedef void* ndr_od_db_handle; //open database handle
//lint -strong(AJX,ndr_od_txn_id)
//object distributor transaction instance identifier
typedef void* ndr_od_txn_id;
#define NDR_OD_INVALID_TXN_ID (ndr_od_txn_id)0
typedef struct //Txn info returned by NdrOdGetTxnInfo
{
ndr_od_db_handle db; /* database */
ndr_dodb_session_type session; /* session */
} ndr_od_txn_info;
//Start a new clash txn for this session
ndr_ret EXPORT
NdrOdStartClashTxn(
ndr_od_db_handle db_handle,
/* -> Handle to the open DB */
ndr_dodb_session_type session, /* -> session */
ndr_od_txn_id *txn_id); /* <- txn id */
//Find out what databases are available
ndr_ret EXPORT
NdrOdEnumerateDBs(
ndr_od_enum_flags flags,
/* -> Determines which databases are included in search*/
ndr_os_db_name search_name,
/* -> The database name (may be wild) */
ndr_os_db_type_name search_type,
/* -> The database type (may be wild) */
ndr_dodb_database_id_type search_id,
/* -> The database id (may be wild) */
ndr_os_db_name name,
/* <- The_database name */
ndr_os_db_type_name type,
/* <- The database type */
ndr_dodb_database_id_type *id,
/* <- The database id */
UINT16 *index);
/* <-> Set to 0 to start else use
previous returned value */
//Start a new txn for this session
ndr_ret EXPORT
NdrOdStartTxn{
ndr_od_db_handle db_handle,
/* -> Handle to the open DB */
ndr_dodb_session_type session,
/* -> session */
ndr_od_txn_id *txn id);
/* -< txn_id */
The interface includes NdrOdcloseTxn(), which closes updates for the current transaction and causes all updates since the last NdrOdStartTxn() call to be applied. Either all updates will be applied, or none will be applied. NdrOdCloseTxn() does not commit the updates, that is, they are not written to disk. NdrOdCommit() is used to commit closed updates to disk. However, after calling NdrOdCloseTxn(), no further updates may be applied in the transaction. This function is also where all the locking and updates previously cached actually get done. Consequently, most locking and/or consistency errors are reported here (after synchronization) so that the transaction can be retried:
ndr_ret EXPORT
NdrOdCloseTxn(ndr_od_txn_id txn_id); /* -> txn_id
*/
The NdrOdEndTxn() function ends the current transaction
and executes an implicit NdrOdCloseTxn(). No error is returned
if no transaction is currently open:
ndr_ret EXPORT
NdrOdEndTxn(ndr_od_txn_id txn_id); /* -> txn id */
The NdrOdCommit function commits all previously closed
transactions for the session to disk:
ndr_ret EXPORT
NdrOdCommit(
ndr_od_db_handle db, /* -> DB to commit */
ndr_dodb_session_type session); /* -> session */
The interface also includes the following functions:
//Abort current txn
ndr_ret EXPORT
NdrOdAbortTxn(ndr_od_txn_id txn_id); /* -> txn_id
*/
//Get info on current txn
ndr_ret EXPORT
NdrOdGetTxnInfo(
ndr_od_txn_id txn id, /* -> txn_id */
ndr_od_txn_info* txn_info); /* <- txn info */
//Lookup an object using parent Distributed Object Identifier
//(DOID; encodes location info to assist in sending distributor
//requests to the right machine; includes UOID) & sibling key
or
//using global key; the key value MUST be a contiguous
structure.
ndr_ret EXPORT
NdrOdLookupByKey(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_os_class class_id,
/* -> Class id. of superclass to match*/
/* Acts as filter when key contains wildcard. */
ndr_dodb_doid_class* parent_doid, /* -> Parent DOID
*/
ndr_os_attribute key_id,
/* -> Type of unique key */
UINT16 key_length,
/* -> Length, in bytes, of the key value */
VOID* key, /* -> Key value
*/
ndr_dodb_doid_class* doid);
/* <- Pointer to returned DOID of object */
//Lookup an object using DOID
//This checks the existence of the object and updates its DOID
ndr_ret EXPORT
NdrOdLookup(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_dodb_doid_class* doid, /* -> DOID */
ndr_dodb_doid_class* new doid);
/* <- Updated DOID of object */
//Lookup an object's parent using DOID.
ndr_ret EXPORT
NdrOdLookupParent(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_dodb_doid_class* doid, /* -> DOID */
ndr_dodb_doid_class* parent_doid);
/* <- Parent DOID of object */
//Read an object using parent DOID and sibling key or using
//global key. It's always OK to read an object with an out of
//date parent doid as the parent's lid is not used to get the
//reference location. The key value MUST be a contiguous
//structure.
ndr_ret EXPORT
NdrOdReadByKey(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_os_Class class_id,
/* -> Class id. of superclass to match */
/* and superclass structure to be returned */
ndr_dodb_doid_class* parent_doid, /* -> Parent DOID */
ndr_os_attribute key_id, /* -> Type of unique key */
UINT16 key_length,
/* -> Length, in bytes, of the key value */
VOID* key, /* -> Key value */
UINT16 max_length,
/* -> Max length of data read */
UINT16* length,
/* <- Final length of data read */
ndr_os_object* object);
/* -> Pointer to object buffer */
//Read an object using DOID
ndr_ret EXPORT
NdrOdRead(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_os_class class_id,
/* -> Class id. of superclass to match */
/* and superclass structure to be returned */
ndr_dodb_doid_class* doid, /* -> DOID */
UINT16 max_length,
/* -> Max length of data read */
UINT16* length,
/* <- Final length of data read */
ndr_os_object* object);
/* -> Pointer to object buffer */
An NdrOdReadn() function which reads multiple objects using parent DOID and wildcards behaves as if none of the updates in the transaction have been applied. Interpretation of wildcard values in the key is done by registered keying functions. NdrOdReadn() reads either up to max_objects, or up to the maximum number of objects that will fit in the max length object buffer:
ndr_ret EXPORT
NdrOdReadn(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_os_class class_id,
/* -> Class id. of superclass to match
and superclass structure to be returned */
ndr_os_class read_as_class,
/* -> Class id. target objects are to be read as */
ndr_dodb_doid_class* parent_doid, /* -> Parent DOID */
ndr_os_attribute key_id, /* -> Type of unique key */
UINT16 key_length,
/* -> Length, in bytes, of the key value */
VOID* key,
/* -> Key value to match, can contain wildcard.
NULL implies match all objects under parent containing
the key id */
UINT16 max_length,
/* -> Max length of data read */
UINT16* length,
/* <- Final length of data read */
ndr_dodb_object_list* object_list,
/* -> Pointer to object buffer */
UINT16 max_objects,
/* -> Max number of objects read. Use OD_MAX_OBJECTS to
read max that will fit in buffer */
ndr_dodb_context_type* context);
/* <> -> set to DODB_CONTEXT_START to start a new read,
or a previously returned context to continue a previous
read. <- set to DODB_CONTEXT_END if all objects read,
or a value that can be used to continue reading at the
next object */
#define NDR_OD_MAX_OBJECTS 0xFFFF
The NdrOdLock() function explicitly adds an exclusive or shared lock to an object using the object's DOID. The lock call is called implicitly for all updates, but should be called explicitly if read locks are required. The lock is only taken when the transaction is initially executed. It is not executed when the update is merged. The lock is applied at the end of a transaction. If it fails the transaction is aborted and should be re-tried by the caller. One embodiment does not utilize locks to control concurrency but instead relies on retries and clash handling:
ndr_ret EXPORT
NdrOdLock(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_doid_class* doid, /* -> Objects's DOID */
BOOLEAN is_exclusive);
/* -> TRUE => take exclusive lock */
The interface also includes:
//Add agent defined lock to object
ndr_ret EXPORT
NdrOdAddAgentLock(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_doid_class* doid, /* -> Objects's DOID */
ndr_dodb_lock_type lock_type,
/* -> Type of lock
ndr_dodb_lock_flags_type lock_flags,
/* -> Flags that allow multiple locks to be taken
in single call. Each bit corresponds to a separate
lock, e.g. used for read/write flags on file open */
ndr_dodb_lock_deny_flags_type deny_flags);
/* -> Bits set that correspond to lock_flags bits
causes the corresponding lock to be denied */
//Remove agent defined lock
ndr_ret EXPORT
NdrOdRemoveAgentLock(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_doid_class* doid, /* -> Objects's DOID */
ndr_dodb_lock_type lock_type);
/* -> Type of lock */
The following four calls are used to append various types of updates onto an open transaction. Any of them may return NDR_OK indicating success, NDR_CD_EXCEEDED_TXN_LIMITS indicating that transaction limits have been exceeded, or some other error indicator. In the case of exceeded transaction limits the transaction state will not have been changed and the failed call will have had no effect. The caller is expected to commit or abort the transaction as appropriate. In all other error cases the transaction is automatically aborted before returning the error to the caller:
//Modify a single attribute in a previously read object
//The object distributor caches the modifications and only
//applies them at close txn time
ndr_ret EXPORT
NdrOdModifyAttribute(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_dodb_doid_class* doid,
/* -> DOID of previous read version of object.
Used to verify object has not been modified by another
user since previously read */
ndr_os_attribute attribute_id,
/* -> Identifies attribute to be modified */
VOID* value); /* -> New attribute value */
//Add a new object
//The DOID attribute does not need to be filled in by the
caller.
//The DOID will be set up before writing the object to the
//database.
ndr_ret EXPORT
NdrOdAdd(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_dodb_doid_class* parent_doid, /* -> Parent DOID */
ndr_os_class class_id,
/* -> Class id of object */
ndr_os_object* object);
/* -> Pointer to agent object */
//Remove an object using DOID
ndr_ret_EXPORT
NdrOdRemove(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_dodb_doid_class* doid); /* -> DOID */
//Move an object using DOID
ndr_ret EXPORT
NdrOdMove(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_access_rights_type rights_needed_on_parent,
/* -> rights needed on parent */
ndr_dodb_doid_class* doid, /* -> DOID */
ndr_dodb_doid_class* target_parent_doid);
/* -> Target parent DOID */
//Set a marker in an open transaction. The state of the
//transaction at the time the marker is set can be reverted
//to at any time before the transaction is closed by
//calling NdrOdRevertToMarker().
//Only the last marker in a transaction is significant.
//This call may return NDR_CD_EXCEEDED_TXN_LIMITS which
//should be treated as for the update appending calls above
ndr_ret EXPORT
NdrOdSetMarker(ndr_od_txn_id txn_id); /* -> txn_id */
//Revert a txn's state to the last previously marked state
ndr_ret EXPORT
NdrOdRevertToMarker(ndr_od_txn_id txn_id); /* -> txn_id */
//Add a <user-id, rights-mask> pair to an object's
//access rights, overwriting any previous rights-mask for
//that user
ndr_ret EXPORT
NdrOdAddAccessRight(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_doid_class* doid, /* -> Object DOID */
ndr_dodb_auth_id_type user,
/* -> User to whom rights are to be granted */
ndr_dodb_access_rights_type rights);
/* -> Rights to be granted to that user */
//Remove any <user-id, rights-mask> pair from an object's
//access rights for a given user-id
ndr_ret EXPORT
NdrOdRemoveAccessRight(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_doid_class* doid, /* -> Object DOID */
ndr_dodb_auth_id_type user);
/* -> User whose rights are to be revoked */
//Get the array of all <user-id, rights-mask> pairs for an
object
ndr_ret EXPORT
NdrOdGetAccessRights(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_doid_class* doid, /* -> Object DOID */
UINT16* acl_count,
/* <- Number of ACL entries for that object */
ndr_dodb_acl_element_type* acl);
/* <- Rights information for that object */
//Get the effective access rights for the current session
//for an object
ndr_ret EXPORT
NdrOdGetEffectiveAccessRight(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_doid_class* doid, /* -> Object DOID */
ndr_dodb_access_rights_type* rights);
/* <- Effective rights for the current session */
//Convert UOID to DOID
ndr_ret EXPORT
NdrOdConvertUoid2Doid(
ndr_os_class class_id,
/* -> Class id. of object */
ndr_dodb_uoid_type* uoid, /* -> UOID */
ndr_dodb_doid_class* doid); /* <- Updated DOID */
//Convert UOID to DOID
ndr_ret EXPORT
NdrOdConvertUoid2LocalDoid(
ndr_os_class class_id,
/* -> Class id. of object */
ndr_dodb_lid_type location,
/* -> Location on which object exists */
ndr_dodb_uoid_type* uoid, /* -> UOID */
ndr_dodb_doid_class* doid); /* <- Updated DOID */
The object processor 86 provides a local hierarchical object-oriented database for objects whose syntax is defined in the object schema 84. In one embodiment, the object processor 86 is built as a layered structure providing functionality according to an interface in the structure which is described below. The embodiment also includes a module for object attribute semantics processing, a set of global secondary indexes, a hierarchy manager, a B-tree manager, a record manager, and a page manager. Suitable modules and managers are readily obtained or constructed by those familiar with database internals. A brief description of the various components follows. The page manager provides functionality according to a logical file interface of free-form fixed length pages addressed by logical page number. Rollback and commit at this level provide anti-crash recovery. The record manager provides for the packing of variable length keyed records into fixed length pages. The B-tree manager uses the facilities of the record and page managers to provide general B-trees supporting variable length records and variable length keys. The hierarchy manager imposes a hierarchical structure on records by use of structured B-tree keys and a global UOID->full name index. The secondary index manager provides generalized global indexing capabilities to records. The attribute manager interprets the schema 84 in order to raise the interface of the object processor 86 from a record-level to an object-level interface. The interface module of the object processor 86 uses lower level interfaces to provide functionality according to the following interface:
op_init Initializes object processor
op_shutdown Shuts down object processor
op_add_database Creates a new volume
op_mount_database Mounts a specified volume for use
op_dismount_database Dismounts the specified volume
op_remove_database Removes a specified volume
(permanently)
op_read Read an object by UOID
op_readn Read one or more objects with
wildcards
op_execute_update_list Apply one or more updates
op_commit Commit updates to a specified
volume
op_rollback Rollback to the last committed
state
op_free_inversion_list Free up an inversion list
returned from update execution
op_clear_stats Clear object processor statistics
op_dump_stats Dump statistics to the log
Due to higher level requirements of trigger functions in a set of trigger function registrations 94, in one embodiment it is necessary to have the old values of modified attributes available on a selective basis. This is done by means of a `preservation list` produced by op_execute_updates(). The preservation list contains an update list specifying old attribute values for all executed updates that require it (as determined by a callback function), together with pointers to the original causative updates. These updates may not actually be present in the input update list, as in the case of an object removal that generates removes for any descendant objects it may have. Preservation lists reside in object processor 86 memory and must thus be freed up by the caller as soon as they are no longer needed. The transaction logger 88 provides a generic transaction log subsystem. The logs maintained by the logger 88 provide keyed access to transaction updates keyed according to location identifier and processor transaction identifier (PTID). In one embodiment, a non-write-through cache is used to batch uncommitted transaction updates. The transaction logger 88 is used by the consistency processor 76 to support fast recovery after a crash. Recovery causes the target database to be updated with any transactions that were committed to the log by the logger 88 but were not written to the target database. The log file header contains a "shutdown OK" flag which is used on startup to determine if recovery is required for the location. The transaction logger 88 is also used by the consistency processor 76 to support fast synchronization. The update log created by the logger 88 is used to replay the updates from one location to a second location using minimal disk and network 10 transfers. The file distributor 90 distributes file contents to appropriate locations in the network 10. A file processor 92 supports each file distributor 90 by carrying out requested read, write, lock, or other operations locally. The file distributor 90 hides from agents the complexities caused by the distributed nature of files. To the extent possible, the interface portion of the file distributor 90 resembles file system interfaces that are familiar in the art. An open file is denoted by a numeric fork_id and functions are provided to read, write, open, and otherwise manipulate and manage files and their contents. However, a class in the schema 84 can be given a REPLICATED_FILE property. Whenever an object of such a class is created in the database, a distributed file is created by the file distributor 90 and file processor 92 to hold the file contents associated with that object. For instance, the Hierarchy Agent might create such an object to denote a leaf node in the directory hierarchy. In short, in one embodiment the file distributor 90 neither has nor needs an explicit externally called mechanism for creating files. Moreover, the distributed file is deleted from storage when the corresponding object is deleted from the database. The locations at which the file is stored are precisely those at which the object exists. When a file with more than one replica 56 is modified and closed, the file distributors 90 and file processors 92 at the various locations holding the replicas 56 ensure that all replicas 56 of the file receive the new contents. It is not necessary for the agent to expressly manage any aspect of file content distribution. A distributed file is identified by the UOID of the corresponding object; no built-in hierarchical naming scheme is used. A transaction identifier is also required when opening a file, to identify the session for which the file is to be opened. In one embodiment, the file distributor 90 and file processor 92 provide functionality according to the following interface:
//An ndr_fd_fork_id is the Id by which an FD open fork is known
typedef SINT16 ndr_fd_fork_id;
#define NDR_FD_NOT_A_FORK_ID (-1)
//An ndr_fd_open_mode is a bit-mask which specifies whether a
//fork is open for reading and/or writing
typedef UINT16 ndr_fd_open_mode;
#define NDR_FD_OPEN_READ_MODE 0 .times. 0001
#define NDR_FD_OPEN_WRITE_MODE 0 .times. 0002
#define NDR_FD_OPEN_EXCL_MODE 0 .times. 0004
#define NDR_FD_OPEN_EXTERNAL_MODES 0 .times. 0007
//The remaming open modes are private to the replica managers
#define NDR_FD_OPEN_SYNC_MODE 0 .times. 0008
#define NDR_FD_OPEN_CLOSE_ON_EOF_MODE 0 .times. 0010
#define NDR_FD_OPEN_READ_NOW 0 .times. 0020
In one alternative embodiment, opening a file with an NdrFdOpenFile() function returns pointers to two functions together with a separate fork_id for use with these two functions only. These pointers are of the type ndr_fd_io_function, and may be used as alternatives to NdrFdReadFile() and NdrFdWriteFile() when accessing that open file only. The functions should be at least as efficient as NdrFdReadFile() and NdrFdWriteFile() and will be significantly faster when the file access is to a local location. Their use does require that the caller maintain a mapping from the open fork id onto these function pointers. For this reason, NdrFdReadFile() and NdrFdWriteFile() should always be available for all open files in this alternative embodiment:
typedef ndr_ret EXPORT (*ndr_fd_io_function)(
ndr_fd_fork_id fork_id, /* -> Id of open fork
*/
UINT32 offset,
/* -> Offset at which to start reading */
UINT16* length,
/* <-> desired length on entry, actual length on
exit. These will only differ if an error
is encountered (such as end of file) */
UINT8* data,
/* <-> Data read or written */
ndr_od_txn_id txn_id); /* -> txn_id */
A "clash" occurs during synchronization when two desired changes to the database are inconsistent. Clashes arise from "independent" updates, namely, updates performed on separate replicas 56 while the computers holding the replicas 56 were disconnected. Thus, clashes always take place between a pair of "clashing updates" which together define a "clash condition." A "repairing update" is an update that removes a clash condition caused by a clashing update. A "transient clash" is a clash that is not present in the final states of the two replicas 56 being merged. Transient clashes only arise when log-based or hybrid merging is used. For instance, suppose two users each create a file of a given name at two locations 36, 38 while those locations are disconnected. The user at the first location 36 then deletes (or renames or moves) the file in question before reconnection such that it no longer clashes with anything on the second location 38. On merging the replicas 56 of the two locations 36, 38, the original add update for the file from the first location 36 will clash with the replica 56 of the second location 38, yet the final result of applying the update stream from the first location 36 to the replica 56 on the second location 38 is a state that is compatible with that replica 56. By contrast, "persistent clashes" create inconsistencies that are present in the final states of two replicas 56. A clash whose type is unknown is a "potential clash." A "file contents clash" occurs when a file's contents have been independently modified on two computers 28, or when a file has been removed from one replica 56 and the file's contents have been independently modified on another replica 56. An "incompatible manipulation clash" occurs when an object's attributes have been independently modified, when an object has been removed in one replica 56 and the object's attributes have been modified in another replica 56, when an object has been removed in one replica 56 and moved in the hierarchy in another replica 56, when a parent object such as a file directory has been removed in one replica 56 and has been given a child object in another replica 56, or when an object has been independently moved in different ways. Thus, although clashes are discussed here in connection with files and the file distributor 90, clashes are not limited to updates involving files. A "unique key clash" occurs when two different objects are given the same key and both objects reside in a portion of the database in which that key should be unique. In a database representing a file system hierarchy, for instance, operations that add, move, or modify files or directories may create a file or directory in one replica 56 that clashes on reconnection with a different but identically-named file or directory in another replica 56. A "permission clash" occurs when a change in file access or modification permissions that is made to a central server replica 56 would prohibit an independent update made to a mobile or client computer replica 56 from being applied to the server replica 56. A permission clash is an example of an "external clash," namely, a clash detected by reference to a structure external to the database. Permission clashes and other external clashes may be detected by trigger functions. A "grouped attribute" is a database object attribute that is associated with other database object attributes such that changing the value of any attribute in a group creates a clash with the other attributes in the group. For instance, filename and rename-inhibit attributes are preferably grouped together, while filename and file-access-date attributes are preferably not grouped together. Without attribute grouping, a change to any attribute of an object is assumed to clash with a change to any other attribute of the object or another change to the same attribute. "Eliminating a clash" means identifying the basis for the clash and eliminating it. "Recovering from a clash" means identifying the basis for the clash and either eliminating that basis or presenting alternative resolutions of the clash to a user to choose from. "Regressing an update" means undoing the update on at least one replica 56. Creating a "recovery item" means creating a duplicate object in a shadow database and then remapping uses of the recovery item's key so that subsequent updates are performed on the recovery item instead of the original object. If the database represents a file system hierarchy, recovery items may be gathered in a "single directory hierarchy" or "recovery directory" that contains a directory at the root of the volume, recovered items, and copies of any directories necessary to connect the recovered items properly with the root. A clash handler function of one of the types below can be registered with the file distributor 90 for a database type to be called whenever the file distributor 90 detects a clash caused by disconnected modification or removal of a file's contents. The parameters are those of a regular clash handler plus the object DOID with NDR_os_CLASS_FLAG_HAS_PARTIALLY_REPLICATED_FILE property (the file object defined by the object schema 84) and possibly a duplicated object return:
//Call back to a husk in respect of clashes detected at the
//database level
typedef ndr_ret EXPORT (*ndr_fd_object_clash_fn)(
ndr_od_db_handle db, /* -> Database */
ndr_dodb_session_type session,
/* -> session to use in od_start_txn */
ndr_od_clash_info* info,
/* -> Information on clash */
ndr_dodb_doid_class* old doid,
/* -> DOID of file with clashing contents */
ndr_dodb_doid_class* new_doid);
/* <- Doid of duplicated file */
//Call back to the husk in respect of clashes detected at the
//filesystem level
// (via pre trigger functions)
typedef ndr_ret EXPORT (*ndr_fd_filesys_clash_fn)(
ndr_od_db_handle db, /* -> Database */
ndr_dodb_session_type session,
/* -> session to use in od_start_txn */
ndr_od_clash_info* info,
/* -> Information on clash */
ndr_dodb_doid_class* doid);
/* -> DOID of file with clashing contents */
A parameter block such as the following is passed to clash handling functions to provide them with information about the clash:
typedef struct
{
ndr_dodb_ptid_type* ptid;
/* -> PTID of clashing txn */
ndr_od_clash_type clash_type;
/* -> Clash type */
ndr_os_class class_id;
/* -> Class id of object causing the clash */
ndr_os_attribute attr_id;
/* -> Attr id of object causing the clash */
ndr_dodb_update_list* update_list;
/* -> Update list of transaction */
ndr_dodb_update* update;
/* -> Update causing clash (always a pointer
into `update_list` */
BOOLEAN is_higher_priority;
/* -> Relative priority of location
to which update is being applied.
TRUE=> Applying to location with higher
priority (e.g. to location set with
central location) */
void* agent_merge_info;
/* -> Value which is reserved for (arbitrary)
use by agent clash handlers. It is
guaranteed to be set to NULL on the
first clash of a merge, and preserved
for all subsequent clashes within that
merge */
} ndr_od_clash_info;
A close handler function of type ndr_fd_close_fn can be registered with the file distributor 90 for a database type to be called whenever the file distributor 90 closes a modified local copy of the file contents, passing the new length and modification date/time and user identifier:
typedef ndr_ret EXPORT (*ndr_fd_close_fn)(
ndr_od_db_handle db, /* -> Database */
ndr_dodb_session_type session,
/* -> session to use in od_start_txn */
ndr_os_class class_id,
/* -> Class ID of file */
ndr_dodb_uoid_type* uoid, /* -> UOID */
UINT32 length,
/* -> length of closed file*/
UINT16 time,
/* -> modification time */
UINT16 date,
/* -> modification date */
UINT32 updator);
/* -> modification user */
A creation handler function of type ndr_fd_creation_fn can be registered with the file distributor 90 for a database type to be called whenever the file distributor 90 creates a local copy of the file contents. This allows the replica manager 46 on a central server computer 28 to update the master copy of the file to reflect the attributes of the file created while disconnected:
typedef ndr_ret EXPORT (*ndr_fd_creation_fn)(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_os_class class_id,
/* -> Class ID of file */
ndr_dodb_uoid_type* uoid); /* -> UOID of file */
The file distributor 90 embodiment also provides the following:
//Return aggregated information about all volumes
ndr_ret EXPORT
NdrFdVolumeInfo(
ndr_od_txn_id txn_id, /* -> txn_id */
UINT32* cluster_size,
/* <- Number of bytes per cluster */
UINT16* total_clusters,
/* <- Total number of clusters */
UINT16* free_clusters);
/* <- Number of free clusters */
//Add a file
ndr_ret EXPORT
NdrFdAddFile(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_doid_class* doid,
/* -> Uoid of file created */
UINT32 length);
/* -> Length of existing file (0 when new) */
//Remove a file
ndr_ret EXPORT
NdrFdRemoveFile(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_uoid_type* uoid);
/* -> Uoid of file removed */
//Open a file for reading or writing by a task
ndr_ret EXPORT
NdrFdOpenFile(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_os_class class_id,
/* -> Class ID of file to open */
ndr_dodb_uoid_type uoid,
/* -> Uoid of file to open */
ndr_fd_open_mode open_mode,
/* -> Open for read and/or write? */
ndr_fd_fork_id* fork_id,
/* <- FD Fork Id of open file */
BOOLEAN is_create,
/* -> TRUE if open as part of create */
ndr_fd_io_function* read_function,
/* <- Function to be used for READ operations */
ndr_fd_io_function* write_function,
/* <- Function to be used for WRITE operations */
ndr_fd_fork_id* io_fork_id,
/* <- FD Fork Id used with above two functions (only) */
UINT16* num_forks_remaining);
/* <- Number of forks remaining to be opened
on same machine */
//Read from a file
ndr_ret EXPORT
NdrFdReadFile(
ndr_od_txn_id txn_id, /* -> txn id */
ndr_fd_fork_id fork_id, /* -> Id of open fork */
UINT32 offset,
/* -> Offset at which to start reading */
UINT16 req_length,
/* -> Number of bytes requested to read */
UINT8* data, /* <- Data read */
UINT16* act_length);
/* <- Actual number of bytes read */
//Write to a file
ndr_ret EXPORT
NdrFdWriteFile(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_fd_fork_id fork_id, /* -> Id of open fork */
UINT32 offset
/* -> Offset at which to start writing */
UINT16 req_length,
/* -> Number of bytes requested to write */
UINT8* data); /* -> Data to be written */
//Get the current length of an open file
ndr_ret EXPORT
NdrFdGetOpenFileLength(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_fd_fork_id fork_id, /* -> Id of open fork */
UINT32* length);
/* <- Length of that open file */
//Lock or Unlock a range of bytes in an open file
ndr_ret EXPORT
NdrFdClearPhysicalRecord ( or NdrFdLockPhysicalRecord(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_fd_fork_id fork_id, /* -> Id of open fork */
UINT32 offset, /* -> Offset for lock */
UINT32 req_length);
/* -> Number of bytes requested to lock */
//Ensure a file's contents are on disk
ndr_ret EXPORT
NdrFdCommitFile(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_fd_fork_id fork_id); /* -> Id of open fork
*/
//Close a file, having completed reading and writing
ndr_ret EXPORT
NdrFdCloseFile(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_fd_fork_id fork_id); /* -> Id of open fork
*/
//Given a UOID to a file or directory return its name
//in the specified namespace, along with its parent's UOID
ndr_ret EXPORT
NdrFdGetFilename(
ndr_od_db_handle db,
/* -> handle to current database */
ndr_dodb_uoid_type* file_or_dir_id,
/* -> Uoid of object whose name is wanted */
ndr_os_attr_property namespace,
/* -> Namespace (e.g. DOS) of name wanted */
void* name_buffer,
/* <- Buffer to receive name */
UINT16* name_size,
/* -> Size of provided buffer */
ndr_dodb_uoid_type* parent_dir_id);
/* <- Parent UOID of object (NULL at root) */
//Callback functions to be used with
//NdrFdRegisterChangedIdCallback
typedef ndr_ret EXPORT
(*NdrFdChangedIdCallback)
ndr_od_db_handle db, /* -> Database Id */
ndr_os_class class_id,
/* -> Class ID of file or dir */
ndr_dodb_uoid_type *uoid, /* -> Uoid of file or dir
*/
UINT32 new_id);
/* -> New Id allocated by underlying file system */
A NdrFdRegisterChangedIdCallback() function provides registration of a callback function to be called when a change to a file or directory's unique identifier is made. On a NetWare 4.x server this normally happens only when the file or directory is created by an internal file distributor 90 trigger function. However the identifier will be needed by agents for tasks such as directory enumeration. Because trigger functions cannot directly modify replicated objects, a record of the identifier change is queued within the file distributor 90 and the callback is made asynchronously:
ndr_ret EXPORT
NdrFdRegisterChangedIdCallback(
ndr_os_db_type_handle db_type, /* -> Database type */
NdrFdChangedIdCallback fn); /* -> Callback function */
The interface also provides the following:
//Register clash handlers for contents clashes for files held
in
//a database of the given type.
ndr_ret EXPORT
NdrFdRegisterClashHandlers(
ndr_os_db_type_handle db_type, // -> Database type
ndr_os_class class_id,
// -> Class ID of contents `container` eg file
ndr_fd_object_clash_fn object_clash_fn,
// -> Clash handler for dealing with conflicts
// -> between objects (e.g. contents modification
// and removal)
ndr_fd_filesys_clash_fn filesys_clash_fn,
// -> Clash handler for conflicts that arise
// through some characteristic of the file
// system (e.g. access rights on delete)
ndr_fd_filesys_clash_fn filesys_clash_fn1);
//Register a trigger-like routine to be called when a local
//replica of a file is modified. The routine takes the length
//and modification date/time of the local replica of the file.
ndr_ret EXPORT
NdrFdRegisterCloseHandler(
ndr_os_db_type_handle db_type, // -> Database type
ndr_os_class class_id,
/* -> Class ID of file */
ndr_fd_close_fn close_fn);
/* -> Clash handler to call */
//Register a trigger-like routine to be called when a local
//replica of a file is has been created. This allows the
//replica manager on a central server to update the
//server's master copy of the file to reflect the attributes
//of the file created during the disconnection.
ndr_ret EXPORT
NdrFdRegisterCreationHandler(
ndr_os_db_type_handle db_type, /* -> Database type */
ndr_os_class class_id,
/* -> Class ID of file */
ndr_fd_creation_fn creation_fn);
/* -> Creation handler to call */
//De-register a clash or close or creation handler for
//contents clashes for files held in a database of the given
type
ndr_ret EXPORT
NdrFdDeRegisterClashHandler( or CloseHandler( or
CreationHandler(
ndr_os_db_type_handle db_type, // -> Database type
ndr_os_class class_id); // -> Class ID of file
//Synchronize all the files to and from this client for the
//passed database. Return control when the files are up to
date.
ndr_ret EXPORT
NdrFdSynchronizeFiles(ndr_od_db_handle db);
//Called from pre trigger functions to check whether
//or not the current connection has sufficient
//per-user-rights to perform a particular operation
//on a particular file system object.
ndr_ret
NdrFdCheckRights(
ndr_dodb_uoid_type* file_uoid,
// uoid of object requiring rights to operation
ndr_od_db_handle db,
// database raising the pre trigger
UINT16 operation);
// bits representing operation
//Note that a file has been locally modified, setting
//modification info and triggering propagation onto other
//replicas.
ndr_ret EXPORT
NdrFdNoteFileModified(
ndr_od_txn_id txn_id, /* -> txn_id */
ndr_dodb_doid_class* file_doid);
The trigger function registrations 94 identify trigger functions that are provided by agents and registered with the object distributor 82. A registered trigger function is called on each event when the associated event occurs. Suitable events include object modification events such as the addition, removal, movement, or modification of an object. Because the trigger functions are called on each location, they can be used to handle mechanisms such as file replication, where the file contents are not stored within the target database, while ensuring that the existence, content, and location of the file tracks the modifications to the target database. All objects must have been locked, either implicitly or via NdrOdLock(), in the triggering transaction before the corresponding trigger function is called, and other objects may only be modified if the trigger function is being called for the first time at the location in question. In an alternative embodiment, the replica manager 46 comprises a NetWare Loadable Module ("NLM") and an NWAdmin snap-in module. The NLM uses hooks in the NetWare file system 48 to intercept updates to the local NetWare storage 54, and uses standard NetWare file system Application Programmer's Interface ("API") calls to update the storage 54 when synchronizing. The architecture is symmetric, with the same code running on all computers 28. The NLM has three major internal subsystems. An environment subsystem provides portability by separating the other two internal subsystems from the operating system environment such as the Windows NT or UNIX environment. The environment subsystem provides execution, debugging, scheduling, thread, and memory management services. A Distributed NetWare ("DNW") subsystem implements NetWare semantics by intercepting NetWare file system calls and calls from a DNW API and making corresponding requests of a dispatch layer discussed below. A distributed responder subsystem implements the replica manager 46 to provide a distributed disconnectable object database which supports replication, transaction synchronization, and schema-definable objects, including file objects, as described herein. An application layer contains application programs and the NWAdmin snap-in. These programs interface with the replica manager 46 either by calling an API or by attempting to access the storage device 54 and being intercepted. An intercept layer in the replica manager 46 intercepts and routes external requests for file system updates that target a replica 56. A dispatch later receives the routed requests and dispatches them to an appropriate agent 44. The agents 44, which have very little knowledge of the distributed nature of the database, invoke the consistency distributor 74, location distributor 78, object distributor 82, and/or file distributor 90. For example, a directory create would result in an object distributor 82 call to NdrOdAdd() to add a new object of type directory. In contrast to the agents 44, the distributors 74,78,82, and 90 have little semantic knowledge of the data but know how it is distributed. The object distributor 82 uses the location distributor 78 to control multi-location operations such as replication and synchronization. The consistency distributor 74 manages transaction semantics, such as when it buffers updates made after a call to NdrOdStartTxn() and applies them atomically when NdrOdEndTxn() is called. The file distributor 90 manages the replication of file contents. The processors 76, 86, 88, and 92 process requests for the local location 40. The consistency processor 76 handles transaction semantics and synchronization, and uses the transaction logger 88 to log updates to the database. The logged updates are used to synchronize other locations 40 and to provide recovery in the event of a clash or a crash. The logger 88 maintains a compressed transaction log. The log is "compressed," for example, in that multiple updates to the "last time modified" attribute of a file object will be represented by a single update. The logger 88 maintains a short sequential on-disk log of recent transactions; the longer-term log is held in the object database as update log entry objects. The object processor 86 implements a local object store and supports the following access methods: hierarchical (e.g., add file object under directory object); global indexed (e.g., read any object using its UOID); and local indexed (e.g., read files and directories within a directory in name order). The object processor 86 uses a variant of a B*-tree. The object processor 86 uses a page table to support atomic commitment of transactional updates, providing rollback and protection against crashes of the computer 40. A file system layer in the file system interface 48 provides a flat file system interface built on the local host file system. It re-maps the flat file system calls to the corresponding files in the hierarchical NetWare volumes to support the current NetWare file system. With reference to FIGS. 1 through 4 and particular focus on FIG. 4, methods of the present invention for managing a transaction log are illustrated. One method of the present invention compresses the transaction log by utilizing an update identifying step 100 to identify a redundant update and a subsequent update removing step 102 to remove that update. These steps, as well as other steps shown in FIG. 4, may be repeated in additional iterations; the bottom of the flowchart leads back to the top as well as to other management operations such as transaction synchronization steps or clash handling steps. A variety of redundant updates are identified by replica managers 46 or other systems which operate according to the methods of the present invention. For instance, a file or directory may be renamed twice, rendering the first rename redundant. Likewise, a file may be modified twice, rendering the first update to the modification date redundant. Scripts or other mechanisms may also repeat operations to no further effect, such as deleting a file twice without recreating it in between the deletes or moving a file and then immediately returning it to its original location. These and similar redundant update sequences are identified during the step 100. More complex but nonetheless redundant sequences can also be analyzed during the step 100. For instance, use of the location state processor 80 may identify an update in the transaction log that specifies an update location on a computer 40 other than the computer 40 which holds the log presently being managed. The log can then be compressed during the step 102 by removing that update. In other situations, further steps are employed to identify redundant updates. For instance, a transaction identifying step 104 determines the most recent successfully merged transaction that updates a selected object. Transaction boundaries may be identified by checkpoints inserted in the log during transaction synchronization or by version control operations. Boundaries may be determined using the object processor 86 as described below in connection with certain three-level structures. Every transaction checkpoint is located at the boundary of a transaction, as defined by the three-level structures or other means, that is found in the log prior to compression. However, not every such boundary will be available as a checkpoint because compression may remove some boundaries. Checkpoints may act as constraints on compression by delimiting incompressible sequences of updates. An update of the selected object that antedates the transaction is next identified during the step 100. Finally, the update is removed during the step 102. This presumes that committed updates will not be needed again. In an alternative embodiment, committed updates are retained to permit recovery from log corruption or to permit the reproduction of earlier versions of objects for other reasons. To rule out certain updates as candidates for removal, an incompressible sequence identifying step 106 is performed in one embodiment. An otherwise removable update will not be removed from an incompressible sequence. A sequence which spans a transaction boundary is incompressible if removing an update from the sequence would require the propagation to replicas 56 of changes to a committed transaction. A sequence which semantically relies on a temporary item to swap two other items cannot be compressed; an example would be the sequence: rename A X rename B A rename X B To facilitate the identifying and/or removing steps 100, 102, the replica manager 46 may reposition an update in the sequence of updates in the transaction log during a repositioning step 108. One or more repositioned updates may then be consolidated with an adjacent update during a consolidating step 110. Consolidation replaces two or more updates with a single equivalent update. For instance, the sequence of updates "add A; add B; add C; rename A Z" may be managed by the repositioning step 108 to produce "add A; rename A Z; add B; add C" and subsequently managed by the consolidating step 110 to produce the sequence "add Z; add B; add C." Of course, other semantically equivalent sequences may also be produced according to the invention, such as the sequence "add B; add C; add Z." In short, redundant updates are identified by examining the operations performed by the updates and the status of the replicas 56. Log compression is based on update semantics, unlike data compression methods such as run-length-encoding which are based only on data values. During a creating step 112, a hierarchical log database is made. The log database represents at least a portion of the transaction log using objects which correspond to the updates and transactions in the specified portion of the transaction log. The log database is preferably efficiently cached and swapped to disk as needed. In one embodiment, the log database is a hierarchical database maintained by the object processor 86. Transactions are represented as a three-level structure. A top-level transaction sequence object contains the PTID associated with the transaction that is described by the object's descendant objects. This PTID is a global key, with its sibling key being the log append order for the transaction in question. An update container object, which is a child of the transaction sequence object, serves as the parent of the transaction's log database update objects. It is separated from the transaction sequence object in order to allow the updates in a transaction to migrate into another (by PTID) transaction during the repositioning step 108. One or more update objects, which are the children of the update container object, represent individual updates to the target database during the transaction. Update objects are separated from one another to aid analysis during the update identifying step 100 and to reduce the effect of possible object size limitations. Update objects are ordered by a sibling key which preserves their order within the update container. In an alternative embodiment, the three-level structure is replaced by a physically flattened, more efficient structure which nonetheless provides logical separation of transaction objects and update objects by use of keys. The namespace of the keys is partitioned to permit interleaving of update and transaction objects under a common parent. The intermediate level, represented in the previously described embodiment by update container objects, is not present. In one embodiment, update compressibility relationships are represented using a "previous update" field in each update object. This field contains the UOID of the update object which represents the most recent previous update in the log for the object in question. A NULL UOID, present only in objects representing a create-performing update, indicates that no such previous update is present. Optional synchronization checkpoints provide a way to group transactions. Checkpoints are represented using multi-valued attributes in transaction sequence objects or one or more checkpoint objects. In one embodiment, one checkpoint object per checkpoint is present; in an alternative embodiment one checkpoint object represents all synchronization checkpoints. Each checkpoint attribute or checkpoint object contains the location identifier value(s) of the location(s) 40 to which the synchronization checkpoint pertains. In the case where the log is on a client 20, this will be the corresponding server 16. If no values are present, no synchronization has yet been done. The portion of the transaction log represented by the log database may be the entire log, or it may be a smaller portion that only includes recent transactions. In the latter case, the remainder of the transaction log is represented by a compressed linearly accessed log stored on the device 54. In embodiments that do not include a log database, the entire transaction log is represented by a linearly accessed log sto | ||||||