-$PostgreSQL: pgsql/src/backend/access/transam/README,v 1.5 2006/03/31 23:32:05 tgl Exp $
+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
In the main processing loop, this results in the following function call
sequence:
- / StartTransactionCommand;
- / StartTransaction;
-1) < ProcessUtility; << BEGIN
- \ BeginTransactionBlock;
- \ CommitTransactionCommand;
+ / StartTransactionCommand;
+ / StartTransaction;
+1) < ProcessUtility; << BEGIN
+ \ BeginTransactionBlock;
+ \ CommitTransactionCommand;
- / StartTransactionCommand;
-2) / ProcessQuery; << SELECT ...
- \ CommitTransactionCommand;
- \ CommandCounterIncrement;
+ / StartTransactionCommand;
+2) / PortalRunSelect; << SELECT ...
+ \ CommitTransactionCommand;
+ \ CommandCounterIncrement;
- / StartTransactionCommand;
-3) / ProcessQuery; << INSERT ...
- \ CommitTransactionCommand;
- \ CommandCounterIncrement;
+ / StartTransactionCommand;
+3) / ProcessQuery; << INSERT ...
+ \ CommitTransactionCommand;
+ \ CommandCounterIncrement;
- / StartTransactionCommand;
- / ProcessUtility; << COMMIT
-4) < EndTransactionBlock;
- \ CommitTransactionCommand;
- \ CommitTransaction;
+ / 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
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
+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
+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)
+ 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
takes care of initializing the memory subsystem at main transaction start.
-Subtransaction handling
+Subtransaction Handling
-----------------------
Subtransactions are implemented using a stack of TransactionState structures,
explicit transaction block has been established, while DefineSavepoint is not.
-Subtransaction numbering
-------------------------
+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.
-A top-level transaction is always given a TransactionId (XID) as soon as it is
-created. This is necessary for a number of reasons, notably XMIN bookkeeping
-for VACUUM. However, a subtransaction doesn't need its own XID unless it
-(or one of its child subxacts) writes tuples into the database. Therefore,
-we postpone assigning XIDs to subxacts until and unless they call
-GetCurrentTransactionId. The subsidiary actions of obtaining a lock on the
-XID and and entering it into pg_subtrans and PG_PROC are done at the same time.
+The subsidiary actions of obtaining a lock on the XID and entering it into
+pg_subtrans and PG_PROC 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 backendID 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 backendID 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
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.)
-
-
-pg_clog and pg_subtrans
+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 MyPgXact->xid 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 MyPgXact, 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 GetOldestXmin, as discussed below.
+
+We allow GetNewTransactionId to store the XID into MyPgXact->xid (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 GetOldestXmin,
+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 MyPgXact->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). GetOldestXmin 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 GetOldestXmin returns. 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 GetOldestXmin'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, GetOldestXmin returns 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.)
+
+GetSnapshotData also performs an oldest-xmin calculation (which had better
+match GetOldestXmin's) and stores that into RecentGlobalXmin, which is used
+for some tuple age cutoff checks where a fresh call of GetOldestXmin seems
+too expensive. Note that while it is certain that two concurrent
+executions of GetSnapshotData will compute the same xmin for their own
+snapshots, as argued above, it is not certain that they will arrive at the
+same estimate of RecentGlobalXmin. This is because we allow XID-less
+transactions to clear their MyPgXact->xmin asynchronously (without taking
+ProcArrayLock), so one execution might see what had been the oldest xmin,
+and another not. This is OK since RecentGlobalXmin 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_clog and pg_subtrans are permanent (on-disk) storage of transaction related
+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_clog records the commit status for each transaction that has been assigned
+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 (either it is
-still running, or the backend crashed without updating its status). A
-sub-committed transaction's status will be updated again to the final value as
-soon as the parent commits or aborts, or when the parent is detected to be
-aborted.
+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
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 the PGPROC struct,
+running --- the main Xid of a transaction is recorded in the PGXACT struct,
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_clog and pg_subtrans. It
+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_clog are implemented in transam.c, while the low-level functions are in
+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
+Write-Ahead Log Coding
----------------------
The WAL subsystem (also called XLOG in the code) exists to guarantee crash
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. Build a WAL log record and pass it to XLogInsert(); then update the page's
-LSN and TLI using the returned XLOG location. For instance,
+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,
- recptr = XLogInsert(rmgr_id, info, rdata);
+ XLogBeginInsert();
+ XLogRegisterBuffer(...)
+ XLogRegisterData(...)
+ recptr = XLogInsert(rmgr_id, info);
PageSetLSN(dp, recptr);
- PageSetTLI(dp, ThisTimeLineID);
6. END_CRIT_SECTION()
7. Unlock and unpin the buffer(s).
-XLogInsert's "rdata" argument is an array of pointer/size items identifying
-chunks of data to be written in the XLOG record, plus optional shared-buffer
-IDs for chunks that are in shared buffers rather than temporary variables.
-The "rdata" array must mention (at least once) each of the shared buffers
-being modified, unless the action is such that the WAL replay routine can
-reconstruct the entire page contents. XLogInsert includes the logic that
-tests to see whether a shared buffer has been modified since the last
-checkpoint. If not, the entire page contents are logged rather than just the
-portion(s) pointed to by "rdata".
-
-Because XLogInsert drops the rdata components associated with buffers it
-chooses to log in full, the WAL replay routines normally need to test to see
-which buffers were handled that way --- otherwise they may be misled about
-what the XLOG record actually contains. XLOG records that describe multi-page
-changes therefore require some care to design: you must be certain that you
-know what data is indicated by each "BKP" bit. An example of the trickiness
-is that in a HEAP_UPDATE record, BKP(1) normally is associated with the source
-page and BKP(2) is associated with the destination page --- but if these are
-the same page, only BKP(1) would have been set.
-
-For this reason as well as the risk of deadlocking on buffer locks, it's best
-to design WAL records so that they reflect small atomic actions involving just
-one or a few pages. The current XLOG infrastructure cannot handle WAL records
-involving references to more than three shared buffers, anyway.
-
-In the case where the WAL record contains enough information to re-generate
-the entire contents of a page, do *not* show that page's buffer ID in the
-rdata array, even if some of the rdata items point into the buffer. This is
-because you don't want XLogInsert to log the whole page contents. The
-standard replay-routine pattern for this case is
-
- reln = XLogOpenRelation(rnode);
- buffer = XLogReadBuffer(reln, blkno, true);
- Assert(BufferIsValid(buffer));
- page = (Page) BufferGetPage(buffer);
-
- ... initialize the page ...
-
- PageSetLSN(page, lsn);
- PageSetTLI(page, ThisTimeLineID);
- MarkBufferDirty(buffer);
- UnlockReleaseBuffer(buffer);
-
-In the case where the WAL record provides only enough information to
-incrementally update the page, the rdata array *must* mention the buffer
-ID at least once; otherwise there is no defense against torn-page problems.
-The standard replay-routine pattern for this case is
-
- if (record->xl_info & XLR_BKP_BLOCK_n)
- << do nothing, page was rewritten from logged copy >>;
-
- reln = XLogOpenRelation(rnode);
- buffer = XLogReadBuffer(reln, blkno, false);
- if (!BufferIsValid(buffer))
- << do nothing, page has been deleted >>;
- page = (Page) BufferGetPage(buffer);
-
- if (XLByteLE(lsn, PageGetLSN(page)))
- {
- /* changes are already applied */
- UnlockReleaseBuffer(buffer);
- return;
- }
-
- ... apply the change ...
-
- PageSetLSN(page, lsn);
- PageSetTLI(page, ThisTimeLineID);
- MarkBufferDirty(buffer);
- UnlockReleaseBuffer(buffer);
-
-As noted above, for a multi-page update you need to be able to determine
-which XLR_BKP_BLOCK_n flag applies to each page. If a WAL record reflects
-a combination of fully-rewritable and incremental updates, then the rewritable
-pages don't count for the XLR_BKP_BLOCK_n numbering. (XLR_BKP_BLOCK_n is
-associated with the n'th distinct buffer ID seen in the "rdata" array, and
-per the above discussion, fully-rewritable buffers shouldn't be mentioned in
-"rdata".)
-
-Due to all these constraints, complex changes (such as a multilevel index
-insertion) normally need to be described by a series of atomic-action WAL
-records. What do you do if the intermediate states are not self-consistent?
-The answer is that the WAL replay logic has to be able to fix things up.
-In btree indexes, for example, a page split requires insertion of a new key in
-the parent btree level, but for locking reasons this has to be reflected by
-two separate WAL records. The replay code has to remember "unfinished" split
-operations, and match them up to subsequent insertions in the parent level.
-If no matching insert has been found by the time the WAL replay ends, the
-replay code has to do the insertion on its own to restore the index to
-consistency.
+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 nrdatas)
+
+ 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 -
+ relfilenode, 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(char *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, char *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 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.
+
+
+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 relfilenode
+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.
+
+
+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 made via any of the paths we have introduced to avoid WAL
+overhead for bulk updates 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, as soon as it finishes
+the bulk update. 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.
projects / postgresql / blobdiff