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src/backend/access/transam/README
913 строк
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Peter Eisentraut
XLogRegisterData, XLogRegisterBufData void * argument for binary data
13 фев 2025, 12:33
13 фев 2025, 12:33
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src/backend/access/transam/README The Transaction System ====================== PostgreSQL's transaction system is a three-layer system. The bottom layer implements low-level transactions and subtransactions, on top of which rests the mainloop's control code, which in turn implements user-visible transactions and savepoints. The middle layer of code is called by postgres.c before and after the processing of each query, or after detecting an error: StartTransactionCommand CommitTransactionCommand AbortCurrentTransaction Meanwhile, the user can alter the system's state by issuing the SQL commands BEGIN, COMMIT, ROLLBACK, SAVEPOINT, ROLLBACK TO or RELEASE. The traffic cop redirects these calls to the toplevel routines BeginTransactionBlock EndTransactionBlock UserAbortTransactionBlock DefineSavepoint RollbackToSavepoint ReleaseSavepoint respectively. Depending on the current state of the system, these functions call low level functions to activate the real transaction system: StartTransaction CommitTransaction AbortTransaction CleanupTransaction StartSubTransaction CommitSubTransaction AbortSubTransaction CleanupSubTransaction Additionally, within a transaction, CommandCounterIncrement is called to increment the command counter, which allows future commands to "see" the effects of previous commands within the same transaction. Note that this is done automatically by CommitTransactionCommand after each query inside a transaction block, but some utility functions also do it internally to allow some operations (usually in the system catalogs) to be seen by future operations in the same utility command. (For example, in DefineRelation it is done after creating the heap so the pg_class row is visible, to be able to lock it.) For example, consider the following sequence of user commands: 1) BEGIN 2) SELECT * FROM foo 3) INSERT INTO foo VALUES (...) 4) COMMIT In the main processing loop, this results in the following function call sequence: / StartTransactionCommand; / StartTransaction; 1) < ProcessUtility; << BEGIN \ BeginTransactionBlock; \ CommitTransactionCommand; / StartTransactionCommand; 2) / PortalRunSelect; << SELECT ... \ CommitTransactionCommand; \ CommandCounterIncrement; / StartTransactionCommand; 3) / ProcessQuery; << INSERT ... \ CommitTransactionCommand; \ CommandCounterIncrement; / StartTransactionCommand; / ProcessUtility; << COMMIT 4) < EndTransactionBlock; \ CommitTransactionCommand; \ CommitTransaction; The point of this example is to demonstrate the need for StartTransactionCommand and CommitTransactionCommand to be state smart -- they should call CommandCounterIncrement between the calls to BeginTransactionBlock and EndTransactionBlock and outside these calls they need to do normal start, commit or abort processing. Furthermore, suppose the "SELECT * FROM foo" caused an abort condition. In this case AbortCurrentTransaction is called, and the transaction is put in aborted state. In this state, any user input is ignored except for transaction-termination statements, or ROLLBACK TO <savepoint> commands. Transaction aborts can occur in two ways: 1) system dies from some internal cause (syntax error, etc) 2) user types ROLLBACK The reason we have to distinguish them is illustrated by the following two situations: case 1 case 2 ------ ------ 1) user types BEGIN 1) user types BEGIN 2) user does something 2) user does something 3) user does not like what 3) system aborts for some reason she sees and types ABORT (syntax error, etc) In case 1, we want to abort the transaction and return to the default state. In case 2, there may be more commands coming our way which are part of the same transaction block; we have to ignore these commands until we see a COMMIT or ROLLBACK. Internal aborts are handled by AbortCurrentTransaction, while user aborts are handled by UserAbortTransactionBlock. Both of them rely on AbortTransaction to do all the real work. The only difference is what state we enter after AbortTransaction does its work: * AbortCurrentTransaction leaves us in TBLOCK_ABORT, * UserAbortTransactionBlock leaves us in TBLOCK_ABORT_END Low-level transaction abort handling is divided in two phases: * AbortTransaction executes as soon as we realize the transaction has failed. It should release all shared resources (locks etc) so that we do not delay other backends unnecessarily. * CleanupTransaction executes when we finally see a user COMMIT or ROLLBACK command; it cleans things up and gets us out of the transaction completely. In particular, we mustn't destroy TopTransactionContext until this point. Also, note that when a transaction is committed, we don't close it right away. Rather it's put in TBLOCK_END state, which means that when CommitTransactionCommand is called after the query has finished processing, the transaction has to be closed. The distinction is subtle but important, because it means that control will leave the xact.c code with the transaction open, and the main loop will be able to keep processing inside the same transaction. So, in a sense, transaction commit is also handled in two phases, the first at EndTransactionBlock and the second at CommitTransactionCommand (which is where CommitTransaction is actually called). The rest of the code in xact.c are routines to support the creation and finishing of transactions and subtransactions. For example, AtStart_Memory takes care of initializing the memory subsystem at main transaction start. Subtransaction Handling ----------------------- Subtransactions are implemented using a stack of TransactionState structures, each of which has a pointer to its parent transaction's struct. When a new subtransaction is to be opened, PushTransaction is called, which creates a new TransactionState, with its parent link pointing to the current transaction. StartSubTransaction is in charge of initializing the new TransactionState to sane values, and properly initializing other subsystems (AtSubStart routines). When closing a subtransaction, either CommitSubTransaction has to be called (if the subtransaction is committing), or AbortSubTransaction and CleanupSubTransaction (if it's aborting). In either case, PopTransaction is called so the system returns to the parent transaction. One important point regarding subtransaction handling is that several may need to be closed in response to a single user command. That's because savepoints have names, and we allow to commit or rollback a savepoint by name, which is not necessarily the one that was last opened. Also a COMMIT or ROLLBACK command must be able to close out the entire stack. We handle this by having the utility command subroutine mark all the state stack entries as commit- pending or abort-pending, and then when the main loop reaches CommitTransactionCommand, the real work is done. The main point of doing things this way is that if we get an error while popping state stack entries, the remaining stack entries still show what we need to do to finish up. In the case of ROLLBACK TO <savepoint>, we abort all the subtransactions up through the one identified by the savepoint name, and then re-create that subtransaction level with the same name. So it's a completely new subtransaction as far as the internals are concerned. Other subsystems are allowed to start "internal" subtransactions, which are handled by BeginInternalSubTransaction. This is to allow implementing exception handling, e.g. in PL/pgSQL. ReleaseCurrentSubTransaction and RollbackAndReleaseCurrentSubTransaction allows the subsystem to close said subtransactions. The main difference between this and the savepoint/release path is that we execute the complete state transition immediately in each subroutine, rather than deferring some work until CommitTransactionCommand. Another difference is that BeginInternalSubTransaction is allowed when no explicit transaction block has been established, while DefineSavepoint is not. Transaction and Subtransaction Numbering ---------------------------------------- Transactions and subtransactions are assigned permanent XIDs only when/if they first do something that requires one --- typically, insert/update/delete a tuple, though there are a few other places that need an XID assigned. If a subtransaction requires an XID, we always first assign one to its parent. This maintains the invariant that child transactions have XIDs later than their parents, which is assumed in a number of places. The subsidiary actions of obtaining a lock on the XID and entering it into pg_subtrans and PGPROC are done at the time it is assigned. A transaction that has no XID still needs to be identified for various purposes, notably holding locks. For this purpose we assign a "virtual transaction ID" or VXID to each top-level transaction. VXIDs are formed from two fields, the procNumber and a backend-local counter; this arrangement allows assignment of a new VXID at transaction start without any contention for shared memory. To ensure that a VXID isn't re-used too soon after backend exit, we store the last local counter value into shared memory at backend exit, and initialize it from the previous value for the same PGPROC slot at backend start. All these counters go back to zero at shared memory re-initialization, but that's OK because VXIDs never appear anywhere on-disk. Internally, a backend needs a way to identify subtransactions whether or not they have XIDs; but this need only lasts as long as the parent top transaction endures. Therefore, we have SubTransactionId, which is somewhat like CommandId in that it's generated from a counter that we reset at the start of each top transaction. The top-level transaction itself has SubTransactionId 1, and subtransactions have IDs 2 and up. (Zero is reserved for InvalidSubTransactionId.) Note that subtransactions do not have their own VXIDs; they use the parent top transaction's VXID. Interlocking Transaction Begin, Transaction End, and Snapshots -------------------------------------------------------------- We try hard to minimize the amount of overhead and lock contention involved in the frequent activities of beginning/ending a transaction and taking a snapshot. Unfortunately, we must have some interlocking for this, because we must ensure consistency about the commit order of transactions. For example, suppose an UPDATE in xact A is blocked by xact B's prior update of the same row, and xact B is doing commit while xact C gets a snapshot. Xact A can complete and commit as soon as B releases its locks. If xact C's GetSnapshotData sees xact B as still running, then it had better see xact A as still running as well, or it will be able to see two tuple versions - one deleted by xact B and one inserted by xact A. Another reason why this would be bad is that C would see (in the row inserted by A) earlier changes by B, and it would be inconsistent for C not to see any of B's changes elsewhere in the database. Formally, the correctness requirement is "if a snapshot A considers transaction X as committed, and any of transaction X's snapshots considered transaction Y as committed, then snapshot A must consider transaction Y as committed". What we actually enforce is strict serialization of commits and rollbacks with snapshot-taking: we do not allow any transaction to exit the set of running transactions while a snapshot is being taken. (This rule is stronger than necessary for consistency, but is relatively simple to enforce, and it assists with some other issues as explained below.) The implementation of this is that GetSnapshotData takes the ProcArrayLock in shared mode (so that multiple backends can take snapshots in parallel), but ProcArrayEndTransaction must take the ProcArrayLock in exclusive mode while clearing the ProcGlobal->xids[] entry at transaction end (either commit or abort). (To reduce context switching, when multiple transactions commit nearly simultaneously, we have one backend take ProcArrayLock and clear the XIDs of multiple processes at once.) ProcArrayEndTransaction also holds the lock while advancing the shared latestCompletedXid variable. This allows GetSnapshotData to use latestCompletedXid + 1 as xmax for its snapshot: there can be no transaction >= this xid value that the snapshot needs to consider as completed. In short, then, the rule is that no transaction may exit the set of currently-running transactions between the time we fetch latestCompletedXid and the time we finish building our snapshot. However, this restriction only applies to transactions that have an XID --- read-only transactions can end without acquiring ProcArrayLock, since they don't affect anyone else's snapshot nor latestCompletedXid. Transaction start, per se, doesn't have any interlocking with these considerations, since we no longer assign an XID immediately at transaction start. But when we do decide to allocate an XID, GetNewTransactionId must store the new XID into the shared ProcArray before releasing XidGenLock. This ensures that all top-level XIDs <= latestCompletedXid are either present in the ProcArray, or not running anymore. (This guarantee doesn't apply to subtransaction XIDs, because of the possibility that there's not room for them in the subxid array; instead we guarantee that they are present or the overflow flag is set.) If a backend released XidGenLock before storing its XID into ProcGlobal->xids[], then it would be possible for another backend to allocate and commit a later XID, causing latestCompletedXid to pass the first backend's XID, before that value became visible in the ProcArray. That would break ComputeXidHorizons, as discussed below. We allow GetNewTransactionId to store the XID into ProcGlobal->xids[] (or the subxid array) without taking ProcArrayLock. This was once necessary to avoid deadlock; while that is no longer the case, it's still beneficial for performance. We are thereby relying on fetch/store of an XID to be atomic, else other backends might see a partially-set XID. This also means that readers of the ProcArray xid fields must be careful to fetch a value only once, rather than assume they can read it multiple times and get the same answer each time. (Use volatile-qualified pointers when doing this, to ensure that the C compiler does exactly what you tell it to.) Another important activity that uses the shared ProcArray is ComputeXidHorizons, which must determine a lower bound for the oldest xmin of any active MVCC snapshot, system-wide. Each individual backend advertises the smallest xmin of its own snapshots in MyProc->xmin, or zero if it currently has no live snapshots (eg, if it's between transactions or hasn't yet set a snapshot for a new transaction). ComputeXidHorizons takes the MIN() of the valid xmin fields. It does this with only shared lock on ProcArrayLock, which means there is a potential race condition against other backends doing GetSnapshotData concurrently: we must be certain that a concurrent backend that is about to set its xmin does not compute an xmin less than what ComputeXidHorizons determines. We ensure that by including all the active XIDs into the MIN() calculation, along with the valid xmins. The rule that transactions can't exit without taking exclusive ProcArrayLock ensures that concurrent holders of shared ProcArrayLock will compute the same minimum of currently-active XIDs: no xact, in particular not the oldest, can exit while we hold shared ProcArrayLock. So ComputeXidHorizons's view of the minimum active XID will be the same as that of any concurrent GetSnapshotData, and so it can't produce an overestimate. If there is no active transaction at all, ComputeXidHorizons uses latestCompletedXid + 1, which is a lower bound for the xmin that might be computed by concurrent or later GetSnapshotData calls. (We know that no XID less than this could be about to appear in the ProcArray, because of the XidGenLock interlock discussed above.) As GetSnapshotData is performance critical, it does not perform an accurate oldest-xmin calculation (it used to, until v14). The contents of a snapshot only depend on the xids of other backends, not their xmin. As backend's xmin changes much more often than its xid, having GetSnapshotData look at xmins can lead to a lot of unnecessary cacheline ping-pong. Instead GetSnapshotData updates approximate thresholds (one that guarantees that all deleted rows older than it can be removed, another determining that deleted rows newer than it can not be removed). GlobalVisTest* uses those thresholds to make invisibility decision, falling back to ComputeXidHorizons if necessary. Note that while it is certain that two concurrent executions of GetSnapshotData will compute the same xmin for their own snapshots, there is no such guarantee for the horizons computed by ComputeXidHorizons. This is because we allow XID-less transactions to clear their MyProc->xmin asynchronously (without taking ProcArrayLock), so one execution might see what had been the oldest xmin, and another not. This is OK since the thresholds need only be a valid lower bound. As noted above, we are already assuming that fetch/store of the xid fields is atomic, so assuming it for xmin as well is no extra risk. pg_xact and pg_subtrans ----------------------- pg_xact and pg_subtrans are permanent (on-disk) storage of transaction related information. There is a limited number of pages of each kept in memory, so in many cases there is no need to actually read from disk. However, if there's a long running transaction or a backend sitting idle with an open transaction, it may be necessary to be able to read and write this information from disk. They also allow information to be permanent across server restarts. pg_xact records the commit status for each transaction that has been assigned an XID. A transaction can be in progress, committed, aborted, or "sub-committed". This last state means that it's a subtransaction that's no longer running, but its parent has not updated its state yet. It is not necessary to update a subtransaction's transaction status to subcommit, so we can just defer it until main transaction commit. The main role of marking transactions as sub-committed is to provide an atomic commit protocol when transaction status is spread across multiple clog pages. As a result, whenever transaction status spreads across multiple pages we must use a two-phase commit protocol: the first phase is to mark the subtransactions as sub-committed, then we mark the top level transaction and all its subtransactions committed (in that order). Thus, subtransactions that have not aborted appear as in-progress even when they have already finished, and the subcommit status appears as a very short transitory state during main transaction commit. Subtransaction abort is always marked in clog as soon as it occurs. When the transaction status all fit in a single CLOG page, we atomically mark them all as committed without bothering with the intermediate sub-commit state. Savepoints are implemented using subtransactions. A subtransaction is a transaction inside a transaction; its commit or abort status is not only dependent on whether it committed itself, but also whether its parent transaction committed. To implement multiple savepoints in a transaction we allow unlimited transaction nesting depth, so any particular subtransaction's commit state is dependent on the commit status of each and every ancestor transaction. The "subtransaction parent" (pg_subtrans) mechanism records, for each transaction with an XID, the TransactionId of its parent transaction. This information is stored as soon as the subtransaction is assigned an XID. Top-level transactions do not have a parent, so they leave their pg_subtrans entries set to the default value of zero (InvalidTransactionId). pg_subtrans is used to check whether the transaction in question is still running --- the main Xid of a transaction is recorded in ProcGlobal->xids[], with a copy in PGPROC->xid, but since we allow arbitrary nesting of subtransactions, we can't fit all Xids in shared memory, so we have to store them on disk. Note, however, that for each transaction we keep a "cache" of Xids that are known to be part of the transaction tree, so we can skip looking at pg_subtrans unless we know the cache has been overflowed. See storage/ipc/procarray.c for the gory details. slru.c is the supporting mechanism for both pg_xact and pg_subtrans. It implements the LRU policy for in-memory buffer pages. The high-level routines for pg_xact are implemented in transam.c, while the low-level functions are in clog.c. pg_subtrans is contained completely in subtrans.c. Write-Ahead Log Coding ---------------------- The WAL subsystem (also called XLOG in the code) exists to guarantee crash recovery. It can also be used to provide point-in-time recovery, as well as hot-standby replication via log shipping. Here are some notes about non-obvious aspects of its design. A basic assumption of a write AHEAD log is that log entries must reach stable storage before the data-page changes they describe. This ensures that replaying the log to its end will bring us to a consistent state where there are no partially-performed transactions. To guarantee this, each data page (either heap or index) is marked with the LSN (log sequence number --- in practice, a WAL file location) of the latest XLOG record affecting the page. Before the bufmgr can write out a dirty page, it must ensure that xlog has been flushed to disk at least up to the page's LSN. This low-level interaction improves performance by not waiting for XLOG I/O until necessary. The LSN check exists only in the shared-buffer manager, not in the local buffer manager used for temp tables; hence operations on temp tables must not be WAL-logged. During WAL replay, we can check the LSN of a page to detect whether the change recorded by the current log entry is already applied (it has been, if the page LSN is >= the log entry's WAL location). Usually, log entries contain just enough information to redo a single incremental update on a page (or small group of pages). This will work only if the filesystem and hardware implement data page writes as atomic actions, so that a page is never left in a corrupt partly-written state. Since that's often an untenable assumption in practice, we log additional information to allow complete reconstruction of modified pages. The first WAL record affecting a given page after a checkpoint is made to contain a copy of the entire page, and we implement replay by restoring that page copy instead of redoing the update. (This is more reliable than the data storage itself would be because we can check the validity of the WAL record's CRC.) We can detect the "first change after checkpoint" by noting whether the page's old LSN precedes the end of WAL as of the last checkpoint (the RedoRecPtr). The general schema for executing a WAL-logged action is 1. Pin and exclusive-lock the shared buffer(s) containing the data page(s) to be modified. 2. START_CRIT_SECTION() (Any error during the next three steps must cause a PANIC because the shared buffers will contain unlogged changes, which we have to ensure don't get to disk. Obviously, you should check conditions such as whether there's enough free space on the page before you start the critical section.) 3. Apply the required changes to the shared buffer(s). 4. Mark the shared buffer(s) as dirty with MarkBufferDirty(). (This must happen before the WAL record is inserted; see notes in SyncOneBuffer().) Note that marking a buffer dirty with MarkBufferDirty() should only happen iff you write a WAL record; see Writing Hints below. 5. If the relation requires WAL-logging, build a WAL record using XLogBeginInsert and XLogRegister* functions, and insert it. (See "Constructing a WAL record" below). Then update the page's LSN using the returned XLOG location. For instance, XLogBeginInsert(); XLogRegisterBuffer(...) XLogRegisterData(...) recptr = XLogInsert(rmgr_id, info); PageSetLSN(dp, recptr); 6. END_CRIT_SECTION() 7. Unlock and unpin the buffer(s). Complex changes (such as a multilevel index insertion) normally need to be described by a series of atomic-action WAL records. The intermediate states must be self-consistent, so that if the replay is interrupted between any two actions, the system is fully functional. In btree indexes, for example, a page split requires a new page to be allocated, and an insertion of a new key in the parent btree level, but for locking reasons this has to be reflected by two separate WAL records. Replaying the first record, to allocate the new page and move tuples to it, sets a flag on the page to indicate that the key has not been inserted to the parent yet. Replaying the second record clears the flag. This intermediate state is never seen by other backends during normal operation, because the lock on the child page is held across the two actions, but will be seen if the operation is interrupted before writing the second WAL record. The search algorithm works with the intermediate state as normal, but if an insertion encounters a page with the incomplete-split flag set, it will finish the interrupted split by inserting the key to the parent, before proceeding. Constructing a WAL record ------------------------- A WAL record consists of a header common to all WAL record types, record-specific data, and information about the data blocks modified. Each modified data block is identified by an ID number, and can optionally have more record-specific data associated with the block. If XLogInsert decides that a full-page image of a block needs to be taken, the data associated with that block is not included. The API for constructing a WAL record consists of five functions: XLogBeginInsert, XLogRegisterBuffer, XLogRegisterData, XLogRegisterBufData, and XLogInsert. First, call XLogBeginInsert(). Then register all the buffers modified, and data needed to replay the changes, using XLogRegister* functions. Finally, insert the constructed record to the WAL by calling XLogInsert(). XLogBeginInsert(); /* register buffers modified as part of this WAL-logged action */ XLogRegisterBuffer(0, lbuffer, REGBUF_STANDARD); XLogRegisterBuffer(1, rbuffer, REGBUF_STANDARD); /* register data that is always included in the WAL record */ XLogRegisterData(&xlrec, SizeOfFictionalAction); /* * register data associated with a buffer. This will not be included * in the record if a full-page image is taken. */ XLogRegisterBufData(0, tuple->data, tuple->len); /* more data associated with the buffer */ XLogRegisterBufData(0, data2, len2); /* * Ok, all the data and buffers to include in the WAL record have * been registered. Insert the record. */ recptr = XLogInsert(RM_FOO_ID, XLOG_FOOBAR_DO_STUFF); Details of the API functions: void XLogBeginInsert(void) Must be called before XLogRegisterBuffer and XLogRegisterData. void XLogResetInsertion(void) Clear any currently registered data and buffers from the WAL record construction workspace. This is only needed if you have already called XLogBeginInsert(), but decide to not insert the record after all. void XLogEnsureRecordSpace(int max_block_id, int ndatas) Normally, the WAL record construction buffers have the following limits: * highest block ID that can be used is 4 (allowing five block references) * Max 20 chunks of registered data These default limits are enough for most record types that change some on-disk structures. For the odd case that requires more data, or needs to modify more buffers, these limits can be raised by calling XLogEnsureRecordSpace(). XLogEnsureRecordSpace() must be called before XLogBeginInsert(), and outside a critical section. void XLogRegisterBuffer(uint8 block_id, Buffer buf, uint8 flags); XLogRegisterBuffer adds information about a data block to the WAL record. block_id is an arbitrary number used to identify this page reference in the redo routine. The information needed to re-find the page at redo - relfilelocator, fork, and block number - are included in the WAL record. XLogInsert will automatically include a full copy of the page contents, if this is the first modification of the buffer since the last checkpoint. It is important to register every buffer modified by the action with XLogRegisterBuffer, to avoid torn-page hazards. The flags control when and how the buffer contents are included in the WAL record. Normally, a full-page image is taken only if the page has not been modified since the last checkpoint, and only if full_page_writes=on or an online backup is in progress. The REGBUF_FORCE_IMAGE flag can be used to force a full-page image to always be included; that is useful e.g. for an operation that rewrites most of the page, so that tracking the details is not worth it. For the rare case where it is not necessary to protect from torn pages, REGBUF_NO_IMAGE flag can be used to suppress full page image from being taken. REGBUF_WILL_INIT also suppresses a full page image, but the redo routine must re-generate the page from scratch, without looking at the old page contents. Re-initializing the page protects from torn page hazards like a full page image does. The REGBUF_STANDARD flag can be specified together with the other flags to indicate that the page follows the standard page layout. It causes the area between pd_lower and pd_upper to be left out from the image, reducing WAL volume. If the REGBUF_KEEP_DATA flag is given, any per-buffer data registered with XLogRegisterBufData() is included in the WAL record even if a full-page image is taken. void XLogRegisterData(const void *data, int len); XLogRegisterData is used to include arbitrary data in the WAL record. If XLogRegisterData() is called multiple times, the data are appended, and will be made available to the redo routine as one contiguous chunk. void XLogRegisterBufData(uint8 block_id, const void *data, int len); XLogRegisterBufData is used to include data associated with a particular buffer that was registered earlier with XLogRegisterBuffer(). If XLogRegisterBufData() is called multiple times with the same block ID, the data are appended, and will be made available to the redo routine as one contiguous chunk. If a full-page image of the buffer is taken at insertion, the data is not included in the WAL record, unless the REGBUF_KEEP_DATA flag is used. Writing a REDO routine ---------------------- A REDO routine uses the data and page references included in the WAL record to reconstruct the new state of the page. The record decoding functions and macros in xlogreader.c/h can be used to extract the data from the record. When replaying a WAL record that describes changes on multiple pages, you must be careful to lock the pages properly to prevent concurrent Hot Standby queries from seeing an inconsistent state. If this requires that two or more buffer locks be held concurrently, you must lock the pages in appropriate order, and not release the locks until all the changes are done. Note that we must only use PageSetLSN/PageGetLSN() when we know the action is serialised. Only Startup process may modify data blocks during recovery, so Startup process may execute PageGetLSN() without fear of serialisation problems. All other processes must only call PageSet/GetLSN when holding either an exclusive buffer lock or a shared lock plus buffer header lock, or be writing the data block directly rather than through shared buffers while holding AccessExclusiveLock on the relation. Writing Hints ------------- In some cases, we write additional information to data blocks without writing a preceding WAL record. This should only happen iff the data can be reconstructed later following a crash and the action is simply a way of optimising for performance. When a hint is written we use MarkBufferDirtyHint() to mark the block dirty. If the buffer is clean and checksums are in use then MarkBufferDirtyHint() inserts an XLOG_FPI_FOR_HINT record to ensure that we take a full page image that includes the hint. We do this to avoid a partial page write, when we write the dirtied page. WAL is not written during recovery, so we simply skip dirtying blocks because of hints when in recovery. If you do decide to optimise away a WAL record, then any calls to MarkBufferDirty() must be replaced by MarkBufferDirtyHint(), otherwise you will expose the risk of partial page writes. The all-visible hint in a heap page (PD_ALL_VISIBLE) is a special case, because it is treated like a durable change in some respects and a hint in other respects. It must satisfy the invariant that, if a heap page's associated visibilitymap (VM) bit is set, then PD_ALL_VISIBLE is set on the heap page itself. Clearing of PD_ALL_VISIBLE is always treated like a fully-durable change to maintain this invariant. Additionally, if checksums or wal_log_hints are enabled, setting PD_ALL_VISIBLE is also treated like a fully-durable change to protect against torn pages. But, if neither checksums nor wal_log_hints are enabled, torn pages are of no consequence if the only change is to PD_ALL_VISIBLE; so no full heap page image is taken, and the heap page's LSN is not updated. NB: it would be incorrect to update the heap page's LSN when applying this optimization, even though there is an associated WAL record, because subsequent modifiers (e.g. an unrelated UPDATE) of the page may falsely believe that a full page image is not required. Write-Ahead Logging for Filesystem Actions ------------------------------------------ The previous section described how to WAL-log actions that only change page contents within shared buffers. For that type of action it is generally possible to check all likely error cases (such as insufficient space on the page) before beginning to make the actual change. Therefore we can make the change and the creation of the associated WAL log record "atomic" by wrapping them into a critical section --- the odds of failure partway through are low enough that PANIC is acceptable if it does happen. Clearly, that approach doesn't work for cases where there's a significant probability of failure within the action to be logged, such as creation of a new file or database. We don't want to PANIC, and we especially don't want to PANIC after having already written a WAL record that says we did the action --- if we did, replay of the record would probably fail again and PANIC again, making the failure unrecoverable. This means that the ordinary WAL rule of "write WAL before the changes it describes" doesn't work, and we need a different design for such cases. There are several basic types of filesystem actions that have this issue. Here is how we deal with each: 1. Adding a disk page to an existing table. This action isn't WAL-logged at all. We extend a table by writing a page of zeroes at its end. We must actually do this write so that we are sure the filesystem has allocated the space. If the write fails we can just error out normally. Once the space is known allocated, we can initialize and fill the page via one or more normal WAL-logged actions. Because it's possible that we crash between extending the file and writing out the WAL entries, we have to treat discovery of an all-zeroes page in a table or index as being a non-error condition. In such cases we can just reclaim the space for re-use. 2. Creating a new table, which requires a new file in the filesystem. We try to create the file, and if successful we make a WAL record saying we did it. If not successful, we can just throw an error. Notice that there is a window where we have created the file but not yet written any WAL about it to disk. If we crash during this window, the file remains on disk as an "orphan". It would be possible to clean up such orphans by having database restart search for files that don't have any committed entry in pg_class, but that currently isn't done because of the possibility of deleting data that is useful for forensic analysis of the crash. Orphan files are harmless --- at worst they waste a bit of disk space --- because we check for on-disk collisions when allocating new relfilenumber OIDs. So cleaning up isn't really necessary. 3. Deleting a table, which requires an unlink() that could fail. Our approach here is to WAL-log the operation first, but to treat failure of the actual unlink() call as a warning rather than error condition. Again, this can leave an orphan file behind, but that's cheap compared to the alternatives. Since we can't actually do the unlink() until after we've committed the DROP TABLE transaction, throwing an error would be out of the question anyway. (It may be worth noting that the WAL entry about the file deletion is actually part of the commit record for the dropping transaction.) 4. Creating and deleting databases and tablespaces, which requires creating and deleting directories and entire directory trees. These cases are handled similarly to creating individual files, ie, we try to do the action first and then write a WAL entry if it succeeded. The potential amount of wasted disk space is rather larger, of course. In the creation case we try to delete the directory tree again if creation fails, so as to reduce the risk of wasted space. Failure partway through a deletion operation results in a corrupt database: the DROP failed, but some of the data is gone anyway. There is little we can do about that, though, and in any case it was presumably data the user no longer wants. In all of these cases, if WAL replay fails to redo the original action we must panic and abort recovery. The DBA will have to manually clean up (for instance, free up some disk space or fix directory permissions) and then restart recovery. This is part of the reason for not writing a WAL entry until we've successfully done the original action. Skipping WAL for New RelFileLocator -------------------------------- Under wal_level=minimal, if a change modifies a relfilenumber that ROLLBACK would unlink, in-tree access methods write no WAL for that change. Code that writes WAL without calling RelationNeedsWAL() must check for this case. This skipping is mandatory. If a WAL-writing change preceded a WAL-skipping change for the same block, REDO could overwrite the WAL-skipping change. If a WAL-writing change followed a WAL-skipping change for the same block, a related problem would arise. When a WAL record contains no full-page image, REDO expects the page to match its contents from just before record insertion. A WAL-skipping change may not reach disk at all, violating REDO's expectation under full_page_writes=off. For any access method, CommitTransaction() writes and fsyncs affected blocks before recording the commit. Prefer to do the same in future access methods. However, two other approaches can work. First, an access method can irreversibly transition a given fork from WAL-skipping to WAL-writing by calling FlushRelationBuffers() and smgrimmedsync(). Second, an access method can opt to write WAL unconditionally for permanent relations. Under these approaches, the access method callbacks must not call functions that react to RelationNeedsWAL(). This applies only to WAL records whose replay would modify bytes stored in the new relfilenumber. It does not apply to other records about the relfilenumber, such as XLOG_SMGR_CREATE. Because it operates at the level of individual relfilenumbers, RelationNeedsWAL() can differ for tightly-coupled relations. Consider "CREATE TABLE t (); BEGIN; ALTER TABLE t ADD c text; ..." in which ALTER TABLE adds a TOAST relation. The TOAST relation will skip WAL, while the table owning it will not. ALTER TABLE SET TABLESPACE will cause a table to skip WAL, but that won't affect its indexes. Asynchronous Commit ------------------- As of PostgreSQL 8.3 it is possible to perform asynchronous commits - i.e., we don't wait while the WAL record for the commit is fsync'ed. We perform an asynchronous commit when synchronous_commit = off. Instead of performing an XLogFlush() up to the LSN of the commit, we merely note the LSN in shared memory. The backend then continues with other work. We record the LSN only for an asynchronous commit, not an abort; there's never any need to flush an abort record, since the presumption after a crash would be that the transaction aborted anyway. We always force synchronous commit when the transaction is deleting relations, to ensure the commit record is down to disk before the relations are removed from the filesystem. Also, certain utility commands that have non-roll-backable side effects (such as filesystem changes) force sync commit to minimize the window in which the filesystem change has been made but the transaction isn't guaranteed committed. The walwriter regularly wakes up (via wal_writer_delay) or is woken up (via its latch, which is set by backends committing asynchronously) and performs an XLogBackgroundFlush(). This checks the location of the last completely filled WAL page. If that has moved forwards, then we write all the changed buffers up to that point, so that under full load we write only whole buffers. If there has been a break in activity and the current WAL page is the same as before, then we find out the LSN of the most recent asynchronous commit, and write up to that point, if required (i.e. if it's in the current WAL page). If more than wal_writer_delay has passed, or more than wal_writer_flush_after blocks have been written, since the last flush, WAL is also flushed up to the current location. This arrangement in itself would guarantee that an async commit record reaches disk after at most two times wal_writer_delay after the transaction completes. However, we also allow XLogFlush to write/flush full buffers "flexibly" (ie, not wrapping around at the end of the circular WAL buffer area), so as to minimize the number of writes issued under high load when multiple WAL pages are filled per walwriter cycle. This makes the worst-case delay three wal_writer_delay cycles. There are some other subtle points to consider with asynchronous commits. First, for each page of CLOG we must remember the LSN of the latest commit affecting the page, so that we can enforce the same flush-WAL-before-write rule that we do for ordinary relation pages. Otherwise the record of the commit might reach disk before the WAL record does. Again, abort records need not factor into this consideration. In fact, we store more than one LSN for each clog page. This relates to the way we set transaction status hint bits during visibility tests. We must not set a transaction-committed hint bit on a relation page and have that record make it to disk prior to the WAL record of the commit. Since visibility tests are normally made while holding buffer share locks, we do not have the option of changing the page's LSN to guarantee WAL synchronization. Instead, we defer the setting of the hint bit if we have not yet flushed WAL as far as the LSN associated with the transaction. This requires tracking the LSN of each unflushed async commit. It is convenient to associate this data with clog buffers: because we will flush WAL before writing a clog page, we know that we do not need to remember a transaction's LSN longer than the clog page holding its commit status remains in memory. However, the naive approach of storing an LSN for each clog position is unattractive: the LSNs are 32x bigger than the two-bit commit status fields, and so we'd need 256K of additional shared memory for each 8K clog buffer page. We choose instead to store a smaller number of LSNs per page, where each LSN is the highest LSN associated with any transaction commit in a contiguous range of transaction IDs on that page. This saves storage at the price of some possibly-unnecessary delay in setting transaction hint bits. How many transactions should share the same cached LSN (N)? If the system's workload consists only of small async-commit transactions, then it's reasonable to have N similar to the number of transactions per walwriter cycle, since that is the granularity with which transactions will become truly committed (and thus hintable) anyway. The worst case is where a sync-commit xact shares a cached LSN with an async-commit xact that commits a bit later; even though we paid to sync the first xact to disk, we won't be able to hint its outputs until the second xact is sync'd, up to three walwriter cycles later. This argues for keeping N (the group size) as small as possible. For the moment we are setting the group size to 32, which makes the LSN cache space the same size as the actual clog buffer space (independently of BLCKSZ). It is useful that we can run both synchronous and asynchronous commit transactions concurrently, but the safety of this is perhaps not immediately obvious. Assume we have two transactions, T1 and T2. The Log Sequence Number (LSN) is the point in the WAL sequence where a transaction commit is recorded, so LSN1 and LSN2 are the commit records of those transactions. If T2 can see changes made by T1 then when T2 commits it must be true that LSN2 follows LSN1. Thus when T2 commits it is certain that all of the changes made by T1 are also now recorded in the WAL. This is true whether T1 was asynchronous or synchronous. As a result, it is safe for asynchronous commits and synchronous commits to work concurrently without endangering data written by synchronous commits. Sub-transactions are not important here since the final write to disk only occurs at the commit of the top level transaction. Changes to data blocks cannot reach disk unless WAL is flushed up to the point of the LSN of the data blocks. Any attempt to write unsafe data to disk will trigger a write which ensures the safety of all data written by that and prior transactions. Data blocks and clog pages are both protected by LSNs. Changes to a temp table are not WAL-logged, hence could reach disk in advance of T1's commit, but we don't care since temp table contents don't survive crashes anyway. Database writes that skip WAL for new relfilenumbers are also safe. In these cases it's entirely possible for the data to reach disk before T1's commit, because T1 will fsync it down to disk without any sort of interlock. However, all these paths are designed to write data that no other transaction can see until after T1 commits. The situation is thus not different from ordinary WAL-logged updates. Transaction Emulation during Recovery ------------------------------------- During Recovery we replay transaction changes in the order they occurred. As part of this replay we emulate some transactional behaviour, so that read only backends can take MVCC snapshots. We do this by maintaining a list of XIDs belonging to transactions that are being replayed, so that each transaction that has recorded WAL records for database writes exist in the array until it commits. Further details are given in comments in procarray.c. Many actions write no WAL records at all, for example read only transactions. These have no effect on MVCC in recovery and we can pretend they never occurred at all. Subtransaction commit does not write a WAL record either and has very little effect, since lock waiters need to wait for the parent transaction to complete. Not all transactional behaviour is emulated, for example we do not insert a transaction entry into the lock table, nor do we maintain the transaction stack in memory. Clog, multixact and commit_ts entries are made normally. Subtrans is maintained during recovery but the details of the transaction tree are ignored and all subtransactions reference the top-level TransactionId directly. Since commit is atomic this provides correct lock wait behaviour yet simplifies emulation of subtransactions considerably. Further details on locking mechanics in recovery are given in comments with the Lock rmgr code.