2003-09-20 22:12:05 +02:00
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<!-- $Header: /cvsroot/pgsql/doc/src/sgml/wal.sgml,v 1.25 2003/09/20 20:12:05 tgl Exp $ -->
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2001-01-25 00:15:19 +01:00
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<chapter id="wal">
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<title>Write-Ahead Logging (<acronym>WAL</acronym>)</title>
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2003-08-31 19:32:24 +02:00
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<indexterm zone="wal">
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<primary>WAL</primary>
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</indexterm>
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<indexterm>
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<primary>transaction log</primary>
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<see>WAL</see>
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</indexterm>
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<para>
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<firstterm>Write-Ahead Logging</firstterm> (<acronym>WAL</acronym>)
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is a standard approach to transaction logging. Its detailed
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description may be found in most (if not all) books about
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transaction processing. Briefly, <acronym>WAL</acronym>'s central
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2001-05-17 23:50:18 +02:00
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concept is that changes to data files (where tables and indexes
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reside) must be written only after those changes have been logged,
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that is, when log records have been flushed to permanent
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storage. If we follow this procedure, we do not need to flush
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data pages to disk on every transaction commit, because we know
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that in the event of a crash we will be able to recover the
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database using the log: any changes that have not been applied to
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the data pages will first be redone from the log records (this is
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roll-forward recovery, also known as REDO) and then changes made by
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uncommitted transactions will be removed from the data pages
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(roll-backward recovery, UNDO).
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</para>
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<sect1 id="wal-benefits-now">
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<title>Benefits of <acronym>WAL</acronym></title>
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<indexterm zone="wal-benefits-now">
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<primary>fsync</primary>
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</indexterm>
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2001-01-25 00:15:19 +01:00
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<para>
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The first obvious benefit of using <acronym>WAL</acronym> is a
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significantly reduced number of disk writes, since only the log
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file needs to be flushed to disk at the time of transaction
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commit; in multiuser environments, commits of many transactions
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may be accomplished with a single <function>fsync()</function> of
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the log file. Furthermore, the log file is written sequentially,
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and so the cost of syncing the log is much less than the cost of
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flushing the data pages.
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</para>
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<para>
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The next benefit is consistency of the data pages. The truth is
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that, before <acronym>WAL</acronym>,
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<productname>PostgreSQL</productname> was never able to guarantee
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consistency in the case of a crash. Before
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<acronym>WAL</acronym>, any crash during writing could result in:
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<orderedlist>
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<listitem>
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<simpara>index rows pointing to nonexistent table rows</simpara>
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</listitem>
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<listitem>
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<simpara>index rows lost in split operations</simpara>
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</listitem>
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<listitem>
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<simpara>totally corrupted table or index page content, because
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of partially written data pages</simpara>
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</listitem>
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</orderedlist>
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2001-05-17 23:50:18 +02:00
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Problems with indexes (problems 1 and 2) could possibly have been
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fixed by additional <function>fsync()</function> calls, but it is
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not obvious how to handle the last case without
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2002-02-12 00:25:14 +01:00
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<acronym>WAL</acronym>; <acronym>WAL</acronym> saves the entire data
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page content in the log if that is required to ensure page
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consistency for after-crash recovery.
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</para>
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</sect1>
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<sect1 id="wal-benefits-later">
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<title>Future Benefits</title>
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<para>
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The UNDO operation is not implemented. This means that changes
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made by aborted transactions will still occupy disk space and that
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a permanent <filename>pg_clog</filename> file to hold
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the status of transactions is still needed, since
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transaction identifiers cannot be reused. Once UNDO is implemented,
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<filename>pg_clog</filename> will no longer be required to be
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permanent; it will be possible to remove
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<filename>pg_clog</filename> at shutdown. (However, the urgency of
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this concern has decreased greatly with the adoption of a segmented
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storage method for <filename>pg_clog</filename>: it is no longer
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necessary to keep old <filename>pg_clog</filename> entries around
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forever.)
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</para>
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<para>
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With UNDO, it will also be possible to implement
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<firstterm>savepoints</firstterm><indexterm><primary>savepoint</></> to allow partial rollback of
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invalid transaction operations (parser errors caused by mistyping
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commands, insertion of duplicate primary/unique keys and so on)
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with the ability to continue or commit valid operations made by
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the transaction before the error. At present, any error will
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invalidate the whole transaction and require a transaction abort.
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</para>
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<para>
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<acronym>WAL</acronym> offers the opportunity for a new method for
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database on-line backup and restore (<acronym>BAR</acronym>). To
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use this method, one would have to make periodic saves of data
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files to another disk, a tape or another host and also archive the
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<acronym>WAL</acronym> log files. The database file copy and the
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archived log files could be used to restore just as if one were
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restoring after a crash. Each time a new database file copy was
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made the old log files could be removed. Implementing this
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facility will require the logging of data file and index creation
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and deletion; it will also require development of a method for
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copying the data files (operating system copy commands are not
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suitable).
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</para>
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<para>
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A difficulty standing in the way of realizing these benefits is that
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they require saving <acronym>WAL</acronym> entries for considerable
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periods of time (e.g., as long as the longest possible transaction if
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transaction UNDO is wanted). The present <acronym>WAL</acronym>
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format is extremely bulky since it includes many disk page
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snapshots. This is not a serious concern at present, since the
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entries only need to be kept for one or two checkpoint intervals;
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but to achieve these future benefits some sort of compressed
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<acronym>WAL</acronym> format will be needed.
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</para>
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</sect1>
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<sect1 id="wal-configuration">
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<title><acronym>WAL</acronym> Configuration</title>
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<para>
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There are several <acronym>WAL</acronym>-related configuration parameters that
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affect database performance. This section explains their use.
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Consult <xref linkend="runtime-config"> for details about setting
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configuration parameters.
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</para>
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<para>
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<firstterm>Checkpoints</firstterm><indexterm><primary>checkpoint</></> are points in the sequence of
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transactions at which it is guaranteed that the data files have
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been updated with all information logged before the checkpoint. At
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checkpoint time, all dirty data pages are flushed to disk and a
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special checkpoint record is written to the log file. As result, in
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the event of a crash, the recoverer knows from what record in the
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log (known as the redo record) it should start the REDO operation,
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since any changes made to data files before that record are already
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on disk. After a checkpoint has been made, any log segments written
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before the redo records are no longer needed and can be recycled or
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removed. (When <acronym>WAL</acronym>-based <acronym>BAR</acronym> is
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implemented, the log segments would be archived before being recycled
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or removed.)
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</para>
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<para>
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The server spawns a special process every so often
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XLOG (and related) changes:
* Store two past checkpoint locations, not just one, in pg_control.
On startup, we fall back to the older checkpoint if the newer one
is unreadable. Also, a physical copy of the newest checkpoint record
is kept in pg_control for possible use in disaster recovery (ie,
complete loss of pg_xlog). Also add a version number for pg_control
itself. Remove archdir from pg_control; it ought to be a GUC
parameter, not a special case (not that it's implemented yet anyway).
* Suppress successive checkpoint records when nothing has been entered
in the WAL log since the last one. This is not so much to avoid I/O
as to make it actually useful to keep track of the last two
checkpoints. If the things are right next to each other then there's
not a lot of redundancy gained...
* Change CRC scheme to a true 64-bit CRC, not a pair of 32-bit CRCs
on alternate bytes. Polynomial borrowed from ECMA DLT1 standard.
* Fix XLOG record length handling so that it will work at BLCKSZ = 32k.
* Change XID allocation to work more like OID allocation. (This is of
dubious necessity, but I think it's a good idea anyway.)
* Fix a number of minor bugs, such as off-by-one logic for XLOG file
wraparound at the 4 gig mark.
* Add documentation and clean up some coding infelicities; move file
format declarations out to include files where planned contrib
utilities can get at them.
* Checkpoint will now occur every CHECKPOINT_SEGMENTS log segments or
every CHECKPOINT_TIMEOUT seconds, whichever comes first. It is also
possible to force a checkpoint by sending SIGUSR1 to the postmaster
(undocumented feature...)
* Defend against kill -9 postmaster by storing shmem block's key and ID
in postmaster.pid lockfile, and checking at startup to ensure that no
processes are still connected to old shmem block (if it still exists).
* Switch backends to accept SIGQUIT rather than SIGUSR1 for emergency
stop, for symmetry with postmaster and xlog utilities. Clean up signal
handling in bootstrap.c so that xlog utilities launched by postmaster
will react to signals better.
* Standalone bootstrap now grabs lockfile in target directory, as added
insurance against running it in parallel with live postmaster.
2001-03-13 02:17:06 +01:00
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to create the next checkpoint. A checkpoint is created every
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<varname>checkpoint_segments</varname> log segments, or every
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<varname>checkpoint_timeout</varname> seconds, whichever comes first.
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XLOG (and related) changes:
* Store two past checkpoint locations, not just one, in pg_control.
On startup, we fall back to the older checkpoint if the newer one
is unreadable. Also, a physical copy of the newest checkpoint record
is kept in pg_control for possible use in disaster recovery (ie,
complete loss of pg_xlog). Also add a version number for pg_control
itself. Remove archdir from pg_control; it ought to be a GUC
parameter, not a special case (not that it's implemented yet anyway).
* Suppress successive checkpoint records when nothing has been entered
in the WAL log since the last one. This is not so much to avoid I/O
as to make it actually useful to keep track of the last two
checkpoints. If the things are right next to each other then there's
not a lot of redundancy gained...
* Change CRC scheme to a true 64-bit CRC, not a pair of 32-bit CRCs
on alternate bytes. Polynomial borrowed from ECMA DLT1 standard.
* Fix XLOG record length handling so that it will work at BLCKSZ = 32k.
* Change XID allocation to work more like OID allocation. (This is of
dubious necessity, but I think it's a good idea anyway.)
* Fix a number of minor bugs, such as off-by-one logic for XLOG file
wraparound at the 4 gig mark.
* Add documentation and clean up some coding infelicities; move file
format declarations out to include files where planned contrib
utilities can get at them.
* Checkpoint will now occur every CHECKPOINT_SEGMENTS log segments or
every CHECKPOINT_TIMEOUT seconds, whichever comes first. It is also
possible to force a checkpoint by sending SIGUSR1 to the postmaster
(undocumented feature...)
* Defend against kill -9 postmaster by storing shmem block's key and ID
in postmaster.pid lockfile, and checking at startup to ensure that no
processes are still connected to old shmem block (if it still exists).
* Switch backends to accept SIGQUIT rather than SIGUSR1 for emergency
stop, for symmetry with postmaster and xlog utilities. Clean up signal
handling in bootstrap.c so that xlog utilities launched by postmaster
will react to signals better.
* Standalone bootstrap now grabs lockfile in target directory, as added
insurance against running it in parallel with live postmaster.
2001-03-13 02:17:06 +01:00
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The default settings are 3 segments and 300 seconds respectively.
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It is also possible to force a checkpoint by using the SQL command
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<command>CHECKPOINT</command>.
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</para>
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2001-10-27 01:10:21 +02:00
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<para>
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Reducing <varname>checkpoint_segments</varname> and/or
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<varname>checkpoint_timeout</varname> causes checkpoints to be done
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more often. This allows faster after-crash recovery (since less work
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will need to be redone). However, one must balance this against the
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increased cost of flushing dirty data pages more often. In addition,
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to ensure data page consistency, the first modification of a data
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page after each checkpoint results in logging the entire page
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content. Thus a smaller checkpoint interval increases the volume of
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output to the log, partially negating the goal of using a smaller
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interval, and in any case causing more disk I/O.
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</para>
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<para>
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There will be at least one 16 MB segment file, and will normally
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not be more than 2 * <varname>checkpoint_segments</varname> + 1
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files. You can use this to estimate space requirements for WAL.
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Ordinarily, when old log segment files are no longer needed, they
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are recycled (renamed to become the next segments in the numbered
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sequence). If, due to a short-term peak of log output rate, there
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are more than 2 * <varname>checkpoint_segments</varname> + 1
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segment files, the unneeded segment files will be deleted instead
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of recycled until the system gets back under this limit.
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</para>
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2002-11-02 23:23:01 +01:00
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<para>
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There are two commonly used <acronym>WAL</acronym> functions:
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<function>LogInsert</function> and <function>LogFlush</function>.
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<function>LogInsert</function> is used to place a new record into
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the <acronym>WAL</acronym> buffers in shared memory. If there is no
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space for the new record, <function>LogInsert</function> will have
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to write (move to kernel cache) a few filled <acronym>WAL</acronym>
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buffers. This is undesirable because <function>LogInsert</function>
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is used on every database low level modification (for example,
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row insertion) at a time when an exclusive lock is held on
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affected data pages, so the operation needs to be as fast as
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possible. What is worse, writing <acronym>WAL</acronym> buffers may
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also force the creation of a new log segment, which takes even more
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time. Normally, <acronym>WAL</acronym> buffers should be written
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and flushed by a <function>LogFlush</function> request, which is
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made, for the most part, at transaction commit time to ensure that
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transaction records are flushed to permanent storage. On systems
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with high log output, <function>LogFlush</function> requests may
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not occur often enough to prevent <acronym>WAL</acronym> buffers
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being written by <function>LogInsert</function>. On such systems
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one should increase the number of <acronym>WAL</acronym> buffers by
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modifying the configuration parameter <varname>wal_buffers</varname>.
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The default number of <acronym>
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WAL</acronym> buffers is 8. Increasing this value will
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correspondingly increase shared memory usage.
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</para>
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2002-11-15 03:44:57 +01:00
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<para>
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Checkpoints are fairly expensive because they force all dirty kernel
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buffers to disk using the operating system <literal>sync()</> call.
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Busy servers may fill checkpoint segment files too quickly,
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causing excessive checkpointing. If such forced checkpoints happen
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more frequently than <varname>checkpoint_warning</varname> seconds,
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a message, will be output to the server logs recommending increasing
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<varname>checkpoint_segments</varname>.
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</para>
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2001-01-25 00:15:19 +01:00
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<para>
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2003-03-24 15:32:51 +01:00
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The <varname>commit_delay</varname> parameter defines for how many
|
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|
microseconds the server process will sleep after writing a commit
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2001-02-18 05:50:43 +01:00
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|
record to the log with <function>LogInsert</function> but before
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2001-01-25 00:15:19 +01:00
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|
performing a <function>LogFlush</function>. This delay allows other
|
2003-03-24 15:32:51 +01:00
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|
server processes to add their commit records to the log so as to have all
|
2003-09-20 22:12:05 +02:00
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|
of them flushed with a single log sync. No sleep will occur if
|
|
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|
<varname>fsync</varname>
|
|
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|
is not enabled, nor if fewer than <varname>commit_siblings</varname>
|
|
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|
other sessions are currently in active transactions; this avoids
|
2003-03-24 15:32:51 +01:00
|
|
|
sleeping when it's unlikely that any other session will commit soon.
|
2001-02-26 01:50:08 +01:00
|
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|
Note that on most platforms, the resolution of a sleep request is
|
2003-03-24 15:32:51 +01:00
|
|
|
ten milliseconds, so that any nonzero <varname>commit_delay</varname>
|
|
|
|
setting between 1 and 10000 microseconds would have the same effect.
|
2001-10-27 01:10:21 +02:00
|
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Good values for these parameters are not yet clear; experimentation
|
|
|
|
is encouraged.
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2001-01-25 00:15:19 +01:00
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</para>
|
2001-03-16 06:44:33 +01:00
|
|
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<para>
|
2003-03-24 15:32:51 +01:00
|
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|
The <varname>wal_sync_method</varname> parameter determines how
|
2001-11-21 07:09:45 +01:00
|
|
|
<productname>PostgreSQL</productname> will ask the kernel to force
|
|
|
|
WAL updates out to disk.
|
2001-03-16 06:44:33 +01:00
|
|
|
All the options should be the same as far as reliability goes,
|
|
|
|
but it's quite platform-specific which one will be the fastest.
|
2003-03-24 15:32:51 +01:00
|
|
|
Note that this parameter is irrelevant if <varname>fsync</varname>
|
2001-03-16 06:44:33 +01:00
|
|
|
has been turned off.
|
|
|
|
</para>
|
|
|
|
|
|
|
|
<para>
|
2003-03-24 15:32:51 +01:00
|
|
|
Setting the <varname>wal_debug</varname> parameter to any nonzero
|
2001-03-16 06:44:33 +01:00
|
|
|
value will result in each <function>LogInsert</function> and
|
|
|
|
<function>LogFlush</function> <acronym>WAL</acronym> call being
|
2003-03-24 15:32:51 +01:00
|
|
|
logged to the server log. At present, it makes no difference what
|
2001-11-28 21:49:10 +01:00
|
|
|
the nonzero value is. This option may be replaced by a more
|
2001-03-16 06:44:33 +01:00
|
|
|
general mechanism in the future.
|
|
|
|
</para>
|
2001-01-25 00:15:19 +01:00
|
|
|
</sect1>
|
2003-03-24 15:32:51 +01:00
|
|
|
|
|
|
|
<sect1 id="wal-internals">
|
|
|
|
<title>Internals</title>
|
|
|
|
|
|
|
|
<para>
|
|
|
|
<acronym>WAL</acronym> is automatically enabled; no action is
|
|
|
|
required from the administrator except ensuring that the additional
|
|
|
|
disk-space requirements of the <acronym>WAL</acronym> logs are met,
|
|
|
|
and that any necessary tuning is done (see <xref
|
|
|
|
linkend="wal-configuration">).
|
|
|
|
</para>
|
|
|
|
|
|
|
|
<para>
|
|
|
|
<acronym>WAL</acronym> logs are stored in the directory
|
|
|
|
<filename>pg_xlog</filename> under the data directory, as a set of
|
|
|
|
segment files, each 16 MB in size. Each segment is divided into 8
|
|
|
|
kB pages. The log record headers are described in
|
|
|
|
<filename>access/xlog.h</filename>; the record content is dependent
|
|
|
|
on the type of event that is being logged. Segment files are given
|
|
|
|
ever-increasing numbers as names, starting at
|
|
|
|
<filename>0000000000000000</filename>. The numbers do not wrap, at
|
|
|
|
present, but it should take a very long time to exhaust the
|
|
|
|
available stock of numbers.
|
|
|
|
</para>
|
|
|
|
|
|
|
|
<para>
|
|
|
|
The <acronym>WAL</acronym> buffers and control structure are in
|
|
|
|
shared memory and are handled by the server child processes; they
|
|
|
|
are protected by lightweight locks. The demand on shared memory is
|
|
|
|
dependent on the number of buffers. The default size of the
|
|
|
|
<acronym>WAL</acronym> buffers is 8 buffers of 8 kB each, or 64 kB
|
|
|
|
total.
|
|
|
|
</para>
|
|
|
|
|
|
|
|
<para>
|
|
|
|
It is of advantage if the log is located on another disk than the
|
|
|
|
main database files. This may be achieved by moving the directory
|
|
|
|
<filename>pg_xlog</filename> to another location (while the server
|
|
|
|
is shut down, of course) and creating a symbolic link from the
|
|
|
|
original location in the main data directory to the new location.
|
|
|
|
</para>
|
|
|
|
|
|
|
|
<para>
|
|
|
|
The aim of <acronym>WAL</acronym>, to ensure that the log is
|
|
|
|
written before database records are altered, may be subverted by
|
2003-08-31 19:32:24 +02:00
|
|
|
disk drives<indexterm><primary>disk drive</></> that falsely report a successful write to the kernel,
|
2003-03-24 15:32:51 +01:00
|
|
|
when, in fact, they have only cached the data and not yet stored it
|
|
|
|
on the disk. A power failure in such a situation may still lead to
|
|
|
|
irrecoverable data corruption. Administrators should try to ensure
|
|
|
|
that disks holding <productname>PostgreSQL</productname>'s
|
|
|
|
<acronym>WAL</acronym> log files do not make such false reports.
|
|
|
|
</para>
|
|
|
|
|
|
|
|
<para>
|
|
|
|
After a checkpoint has been made and the log flushed, the
|
|
|
|
checkpoint's position is saved in the file
|
|
|
|
<filename>pg_control</filename>. Therefore, when recovery is to be
|
|
|
|
done, the server first reads <filename>pg_control</filename> and
|
|
|
|
then the checkpoint record; then it performs the REDO operation by
|
|
|
|
scanning forward from the log position indicated in the checkpoint
|
|
|
|
record. Because the entire content of data pages is saved in the
|
|
|
|
log on the first page modification after a checkpoint, all pages
|
|
|
|
changed since the checkpoint will be restored to a consistent
|
|
|
|
state.
|
|
|
|
</para>
|
|
|
|
|
|
|
|
<para>
|
|
|
|
Using <filename>pg_control</filename> to get the checkpoint
|
|
|
|
position speeds up the recovery process, but to handle possible
|
|
|
|
corruption of <filename>pg_control</filename>, we should actually
|
|
|
|
implement the reading of existing log segments in reverse order --
|
|
|
|
newest to oldest -- in order to find the last checkpoint. This has
|
|
|
|
not been implemented, yet.
|
|
|
|
</para>
|
|
|
|
</sect1>
|
2001-01-25 00:15:19 +01:00
|
|
|
</chapter>
|
|
|
|
|
|
|
|
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