The Sparc machines in the buildfarm are crashing because of misaligned
access to posting lists stored in entry tuples.
I accidentally removed a critical SHORTALIGN() from ginFormTuple, as part
of the packed posting lists patch. Perhaps I thought it was unnecessary,
because the index_form_tuple() call above the SHORTALIGN already aligned
the size, missing the fact that the null-category byte makes it misaligned
again (I think the SHORTALIGN is indeed unnecessary if there's no null-
category byte, but let's just play it safe...)
GIN posting lists are now encoded using varbyte-encoding, which allows them
to fit in much smaller space than the straight ItemPointer array format used
before. The new encoding is used for both the lists stored in-line in entry
tree items, and in posting tree leaf pages.
To maintain backwards-compatibility and keep pg_upgrade working, the code
can still read old-style pages and tuples. Posting tree leaf pages in the
new format are flagged with GIN_COMPRESSED flag, to distinguish old and new
format pages. Likewise, entry tree tuples in the new format have a
GIN_ITUP_COMPRESSED flag set in a bit that was previously unused.
This patch bumps GIN_CURRENT_VERSION from 1 to 2. New indexes created with
version 9.4 will therefore have version number 2 in the metapage, while old
pg_upgraded indexes will have version 1. The code treats them the same, but
it might be come handy in the future, if we want to drop support for the
uncompressed format.
Alexander Korotkov and me. Reviewed by Tomas Vondra and Amit Langote.
In ordinary operation, VACUUM must be careful to take a cleanup lock on
each leaf page of a btree index; this ensures that no indexscans could
still be "in flight" to heap tuples due to be deleted. (Because of
possible index-tuple motion due to concurrent page splits, it's not enough
to lock only the pages we're deleting index tuples from.) In Hot Standby,
the WAL replay process must likewise lock every leaf page. There were
several bugs in the code for that:
* The replay scan might come across unused, all-zero pages in the index.
While btree_xlog_vacuum itself did the right thing (ie, nothing) with
such pages, xlogutils.c supposed that such pages must be corrupt and
would throw an error. This accounts for various reports of replication
failures with "PANIC: WAL contains references to invalid pages". To
fix, add a ReadBufferMode value that instructs XLogReadBufferExtended
not to complain when we're doing this.
* btree_xlog_vacuum performed the extra locking if standbyState ==
STANDBY_SNAPSHOT_READY, but that's not the correct test: we won't open up
for hot standby queries until the database has reached consistency, and
we don't want to do the extra locking till then either, for fear of reading
corrupted pages (which bufmgr.c would complain about). Fix by exporting a
new function from xlog.c that will report whether we're actually in hot
standby replay mode.
* To ensure full coverage of the index in the replay scan, btvacuumscan
would emit a dummy WAL record for the last page of the index, if no
vacuuming work had been done on that page. However, if the last page
of the index is all-zero, that would result in corruption of said page,
since the functions called on it weren't prepared to handle that case.
There's no need to lock any such pages, so change the logic to target
the last normal leaf page instead.
The first two of these bugs were diagnosed by Andres Freund, the other one
by me. Fixes based on ideas from Heikki Linnakangas and myself.
This has been wrong since Hot Standby was introduced, so back-patch to 9.0.
The new MultiXact freezing routines introduced by commit 8e9a16ab8f
neglected to consider tuples that came from a pg_upgrade'd database; a
vacuum run that tried to freeze such tuples would die with an error such
as
ERROR: MultiXactId 11415437 does no longer exist -- apparent wraparound
To fix, ensure that GetMultiXactIdMembers is allowed to return empty
multis when the infomask bits are right, as is done in other callsites.
Per trouble report from F-Secure.
In passing, fix a copy&paste bug reported by Andrey Karpov from VIVA64
from their PVS-Studio static checked, that instead of setting relminmxid
to Invalid, we were setting relfrozenxid twice. Not an important
mistake because that code branch is about relations for which we don't
use the frozenxid/minmxid values at all in the first place, but seems to
warrants a fix nonetheless.
Makes the replay loop slightly more readable, by separating the concerns of
whether to stop and whether to delay, and how to extract the timestamp from
a record.
This has the user-visible change that the timestamp of the last applied
record is now updated after actually applying it. Before, it was updated
just before applying it. That meant that pg_last_xact_replay_timestamp()
could return the timestamp of a commit record that is in process of being
replayed, but not yet applied. Normally the difference is small, but if
min_recovery_apply_delay is set, there could be a significant delay between
reading a record and applying it.
Another behavioral change is that if you recover to a restore point, we stop
after the restore point record, not before it. It makes no difference as far
as running queries on the server is concerned, as applying a restore point
record changes nothing, but if examine the timeline history you will see
that the new timeline branched off just after the restore point record, not
before it. One practical consequence is that if you do PITR to the new
timeline, and set recovery target to the same named restore point again, it
will find and stop recovery at the same restore point. Conceptually, I think
it makes more sense to consider the restore point as part of the new
timeline's history than not.
In principle, setting the last-replayed timestamp before actually applying
the record was a bug all along, but it doesn't seem worth the risk to
backpatch, since min_recovery_apply_delay was only added in 9.4.
If pause_at_recovery_target is set, recovery pauses *before* applying the
target record, even if recovery_target_inclusive is set. If you then
continue with pg_xlog_replay_resume(), it will apply the target record
before ending recovery. In other words, if you log in while it's paused
and verify that the database looks OK, ending recovery changes its state
again, possibly destroying data that you were tring to salvage with PITR.
Backpatch to 9.1, this has been broken since pause_at_recovery_target was
added.
The docs say that only one of recovery_target_xid, recovery_target_time, or
recovery_target_name can be specified. But the code actually did something
different, so that a name overrode time, and xid overrode both time and name.
Now the target specified last takes effect, whether it's an xid, time or
name.
With this patch, we still accept multiple recovery_target settings, even
though docs say that only one can be specified. It's a general property of
the recovery.conf file parser that you if you specify the same option twice,
the last one takes effect, like with postgresql.conf.
When starting WAL replay from an online checkpoint, the last replayed WAL
record variable was initialized using the checkpoint record's location, even
though the records between the REDO location and the checkpoint record had
not been replayed yet. That was noted as "slightly confusing" but harmless
in the comment, but in some cases, it fooled CheckRecoveryConsistency to
incorrectly conclude that we had already reached a consistent state
immediately at the beginning of WAL replay. That caused the system to accept
read-only connections in hot standby mode too early, and also PANICs with
message "WAL contains references to invalid pages".
Fix by initializing the variables to the REDO location instead.
In 9.2 and above, change CheckRecoveryConsistency() to use
lastReplayedEndRecPtr variable when checking if backup end location has
been reached. It was inconsistently using EndRecPtr for that check, but
lastReplayedEndRecPtr when checking min recovery point. It made no
difference before this patch, because in all the places where
CheckRecoveryConsistency was called the two variables were the same, but
it was always an accident waiting to happen, and would have been wrong
after this patch anyway.
Report and analysis by Tomonari Katsumata, bug #8686. Backpatch to 9.0,
where hot standby was introduced.
And the same for do_pg_stop_backup. The code in do_pg_* is not allowed
to access the catalogs. For manual base backups, the permissions
check can be handled in the calling function, and for streaming
base backups only users with the required permissions can get past
the authentication step in the first place.
Reported by Antonin Houska, diagnosed by Andres Freund
Original users of slru.c were all producing 4-digit filenames, so that
was all that that code was prepared to handle. Changes to multixact.c
in the course of commit 0ac5ad5134 made pg_multixact/members create
5-digit filenames once a certain threshold was reached, which
SlruScanDirectory wasn't prepared to deal with; in particular,
5-digit-name files were not removed during truncation. Change that
routine to make it aware of those files, and have it process them just
like any others.
Right now, some pg_multixact/members directories will contain a mixture
of 4-char and 5-char filenames. A future commit is expected fix things
so that each slru.c user declares the correct maximum width for the
files it produces, to avoid such unsightly mixtures.
Noticed while investigating bug #8673 reported by Serge Negodyuck.
In the 9.2 code for extending multixact/members, the logic was very
simple because the number of entries in a members page was a proper
divisor of 2^32, and thus at 2^32 wraparound the logic for page switch
was identical than at any other page boundary. In commit 0ac5ad5134 I
failed to realize this and introduced code that was not able to go over
the 2^32 boundary. Fix that by ensuring that when we reach the last
page of the last segment we correctly zero the initial page of the
initial segment, using correct uint32-wraparound-safe arithmetic.
Noticed while investigating bug #8673 reported by Serge Negodyuck, as
diagnosed by Andres Freund.
In pg_multixact/members, relying on modulo-2^32 arithmetic for
wraparound handling doesn't work all that well. Because we don't
explicitely track wraparound of the allocation counter for members, it
is possible that the "live" area exceeds 2^31 entries; trying to remove
SLRU segments that are "old" according to the original logic might lead
to removal of segments still in use. To fix, have the truncation
routine use a tailored SlruScanDirectory callback that keeps track of
the live area in actual use; that way, when the live range exceeds 2^31
entries, the oldest segments still live will not get removed untimely.
This new SlruScanDir callback needs to take care not to remove segments
that are "in the future": if new SLRU segments appear while the
truncation is ongoing, make sure we don't remove them. This requires
examination of shared memory state to recheck for false positives, but
testing suggests that this doesn't cause a problem. The original coding
didn't suffer from this pitfall because segments created when truncation
is running are never considered to be removable.
Per Andres Freund's investigation of bug #8673 reported by Serge
Negodyuck.
We haven't wanted to do this in the past on the grounds that in rare
cases the original xmin value will be needed for forensic purposes, but
commit 37484ad2aa removes that objection,
so now we can.
Per extensive discussion, among many people, on pgsql-hackers.
Commit 37484ad2aa invalidated a good
chunk of documentation, so patch it up to reflect the new state of
play. Along the way, patch remaining documentation references to
FrozenXID to say instead FrozenTransactionId, so that they match the
way we actually spell it in the code.
Instead of changing the tuple xmin to FrozenTransactionId, the combination
of HEAP_XMIN_COMMITTED and HEAP_XMIN_INVALID, which were previously never
set together, is now defined as HEAP_XMIN_FROZEN. A variety of previous
proposals to freeze tuples opportunistically before vacuum_freeze_min_age
is reached have foundered on the objection that replacing xmin by
FrozenTransactionId might hinder debugging efforts when things in this
area go awry; this patch is intended to solve that problem by keeping
the XID around (but largely ignoring the value to which it is set).
Third-party code that checks for HEAP_XMIN_INVALID on tuples where
HEAP_XMIN_COMMITTED might be set will be broken by this change. To fix,
use the new accessor macros in htup_details.h rather than consulting the
bits directly. HeapTupleHeaderGetXmin has been modified to return
FrozenTransactionId when the infomask bits indicate that the tuple is
frozen; use HeapTupleHeaderGetRawXmin when you already know that the
tuple isn't marked commited or frozen, or want the raw value anyway.
We currently do this in routines that display the xmin for user consumption,
in tqual.c where it's known to be safe and important for the avoidance of
extra cycles, and in the function-caching code for various procedural
languages, which shouldn't invalidate the cache just because the tuple
gets frozen.
Robert Haas and Andres Freund
Updating or locking a row that was already locked by the same
transaction under the same Xid caused a MultiXact to be created; but
this is unnecessary, because there's no usefulness in being able to
differentiate two locks by the same transaction. In particular, if a
transaction executed SELECT FOR UPDATE followed by an UPDATE that didn't
modify columns of the key, we would dutifully represent the resulting
combination as a multixact -- even though a single key-update is
sufficient.
Optimize the case so that only the strongest of both locks/updates is
represented in Xmax. This can save some Xmax's from becoming
MultiXacts, which can be a significant optimization.
This missed optimization opportunity was spotted by Andres Freund while
investigating a bug reported by Oliver Seemann in message
CANCipfpfzoYnOz5jj=UZ70_R=CwDHv36dqWSpwsi27vpm1z5sA@mail.gmail.com
and also directly as a performance regression reported by Dong Ye in
message
d54b8387.000012d8.00000010@YED-DEVD1.vmware.com
Reportedly, this patch fixes the performance regression.
Since the missing optimization was reported as a significant performance
regression from 9.2, backpatch to 9.3.
Andres Freund, tweaked by Álvaro Herrera
If a tuple was locked by transaction A, and transaction B updated it,
the new version of the tuple created by B would be locked by A, yet
visible only to B; due to an oversight in HeapTupleSatisfiesUpdate, the
lock held by A wouldn't get checked if transaction B later deleted (or
key-updated) the new version of the tuple. This might cause referential
integrity checks to give false positives (that is, allow deletes that
should have been rejected).
This is an easy oversight to have made, because prior to improved tuple
locks in commit 0ac5ad5134 it wasn't possible to have tuples created by
our own transaction that were also locked by remote transactions, and so
locks weren't even considered in that code path.
It is recommended that foreign keys be rechecked manually in bulk after
installing this update, in case some referenced rows are missing with
some referencing row remaining.
Per bug reported by Daniel Wood in
CAPweHKe5QQ1747X2c0tA=5zf4YnS2xcvGf13Opd-1Mq24rF1cQ@mail.gmail.com
Tuple freezing was broken in connection to MultiXactIds; commit
8e53ae025d tried to fix it, but didn't go far enough. As noted by
Noah Misch, freezing a tuple whose Xmax is a multi containing an aborted
update might cause locks in the multi to go ignored by later
transactions. This is because the code depended on a multixact above
their cutoff point not having any lock-only member older than the cutoff
point for Xids, which is easily defeated in READ COMMITTED transactions.
The fix for this involves creating a new MultiXactId when necessary.
But this cannot be done during WAL replay, and moreover multixact
examination requires using CLOG access routines which are not supposed
to be used during WAL replay either; so tuple freezing cannot be done
with the old freeze WAL record. Therefore, separate the freezing
computation from its execution, and change the WAL record to carry all
necessary information. At WAL replay time, it's easy to re-execute
freezing because we don't need to re-compute the new infomask/Xmax
values but just take them from the WAL record.
While at it, restructure the coding to ensure all page changes occur in
a single critical section without much room for failures. The previous
coding wasn't using a critical section, without any explanation as to
why this was acceptable.
In replication scenarios using the 9.3 branch, standby servers must be
upgraded before their master, so that they are prepared to deal with the
new WAL record once the master is upgraded; failure to do so will cause
WAL replay to die with a PANIC message. Later upgrade of the standby
will allow the process to continue where it left off, so there's no
disruption of the data in the standby in any case. Standbys know how to
deal with the old WAL record, so it's okay to keep the master running
the old code for a while.
In master, the old freeze WAL record is gone, for cleanliness' sake;
there's no compatibility concern there.
Backpatch to 9.3, where the original bug was introduced and where the
previous fix was backpatched.
Álvaro Herrera and Andres Freund
The original performs too poorly; in some scenarios it shows way too
high while profiling. Try to make it a bit smarter to avoid excessive
cosst. In particular, make it have a maximum size, and have entries be
sorted in LRU order; once the max size is reached, evict the oldest
entry to avoid it from growing too large.
Per complaint from Andres Freund in connection with new tuple freezing
code.
WAL records of hint bit updates is useful to tools that want to examine
which pages have been modified. In particular, this is required to make
the pg_rewind tool safe (without checksums).
This can also be used to test how much extra WAL-logging would occur if
you enabled checksums, without actually enabling them (which you can't
currently do without re-initdb'ing).
Sawada Masahiko, docs by Samrat Revagade. Reviewed by Dilip Kumar, with
further changes by me.
Set min_recovery_apply_delay to force a delay in recovery apply for commit and
restore point WAL records. Other records are replayed immediately. Delay is
measured between WAL record time and local standby time.
Robert Haas, Fabrízio de Royes Mello and Simon Riggs
Detailed review by Mitsumasa Kondo
There's no real point in not doing this. It doesn't cost anything
in performance or space. So let's go wild.
Andres Freund, with substantial editing as to style by me.
When this reloption is set and wal_level=logical is configured,
we'll record the CIDs stamped by inserts, updates, and deletes to
the table just as we would for an actual catalog table. This will
allow logical decoding to use historical MVCC snapshots to access
such tables just as they access ordinary catalog tables.
Replication solutions built around the logical decoding machinery
will likely need to set this operation for their configuration
tables; it might also be needed by extensions which perform table
access in their output functions.
Andres Freund, reviewed by myself and others.
When wal_level=logical, we'll log columns from the old tuple as
configured by the REPLICA IDENTITY facility added in commit
07cacba983. This makes it possible
a properly-configured logical replication solution to correctly
follow table updates even if they change the chosen key columns,
or, with REPLICA IDENTITY FULL, even if the table has no key at
all. Note that updates which do not modify the replica identity
column won't log anything extra, making the choice of a good key
(i.e. one that will rarely be changed) important to performance
when wal_level=logical is configured.
Each insert, update, or delete to a catalog table will also log
the CMIN and/or CMAX values of stamped by the current transaction.
This is necessary because logical decoding will require access to
historical snapshots of the catalog in order to decode some data
types, and the CMIN/CMAX values that we may need in order to judge
row visibility may have been overwritten by the time we need them.
Andres Freund, reviewed in various versions by myself, Heikki
Linnakangas, KONDO Mitsumasa, and many others.
In 247c76a989, I added some code to do fine-grained checking of
MultiXact status of locking/updating transactions when traversing an
update chain. There was a thinko in that patch which would have the
traversing abort, that is return HeapTupleUpdated, when the other
transaction is a committed lock-only. In this case we should ignore it
and return success instead. Of course, in the case where there is a
committed update, HeapTupleUpdated is the correct return value.
A user-visible symptom of this bug is that in REPEATABLE READ and
SERIALIZABLE transaction isolation modes spurious serializability errors
can occur:
ERROR: could not serialize access due to concurrent update
In order for this to happen, there needs to be a tuple that's key-share-
locked and also updated, and the update must abort; a subsequent
transaction trying to acquire a new lock on that tuple would abort with
the above error. The reason is that the initial FOR KEY SHARE is seen
as committed by the new locking transaction, which triggers this bug.
(If the UPDATE commits, then the serialization error is correctly
reported.)
When running a query in READ COMMITTED mode, what happens is that the
locking is aborted by the HeapTupleUpdated return value, then
EvalPlanQual fetches the newest version of the tuple, which is then the
only version that gets locked. (The second time the tuple is checked
there is no misbehavior on the committed lock-only, because it's not
checked by the code that traverses update chains; so no bug.) Only the
newest version of the tuple is locked, not older ones, but this is
harmless.
The isolation test added by this commit illustrates the desired
behavior, including the proper serialization errors that get thrown.
Backpatch to 9.3.
Insertion to a non-leaf GIN page didn't make a full-page image of the page,
which is wrong. The code used to do it correctly, but was changed (commit
853d1c3103) because the redo-routine didn't
track incomplete splits correctly when the page was restored from a full
page image. Of course, that was not right way to fix it, the redo routine
should've been fixed instead. The redo-routine was surreptitiously fixed
in 2010 (commit 4016bdef8a), so all we need
to do now is revert the code that creates the record to its original form.
This doesn't change the format of the WAL record.
Backpatch to all supported versions.
When an external recovery command such as restore_command or
archive_cleanup_command fails, report the exit code properly,
distinguishing signals and normal exists, using the existing
wait_result_to_str() facility, instead of just reporting the return
value from system().
Reviewed-by: Peter Geoghegan <pg@heroku.com>
Both heap_freeze_tuple() and heap_tuple_needs_freeze() neglected to look
into a multixact to check the members against cutoff_xid. This means
that a very old Xid could survive hidden within a multi, possibly
outliving its CLOG storage. In the distant future, this would cause
clog lookup failures:
ERROR: could not access status of transaction 3883960912
DETAIL: Could not open file "pg_clog/0E78": No such file or directory.
This mostly was problematic when the updating transaction aborted, since
in that case the row wouldn't get pruned away earlier in vacuum and the
multixact could possibly survive for a long time. In many cases, data
that is inaccessible for this reason way can be brought back
heuristically.
As a second bug, heap_freeze_tuple() didn't properly handle multixacts
that need to be frozen according to cutoff_multi, but whose updater xid
is still alive. Instead of preserving the update Xid, it just set Xmax
invalid, which leads to both old and new tuple versions becoming
visible. This is pretty rare in practice, but a real threat
nonetheless. Existing corrupted rows, unfortunately, cannot be repaired
in an automated fashion.
Existing physical replicas might have already incorrectly frozen tuples
because of different behavior than in master, which might only become
apparent in the future once pg_multixact/ is truncated; it is
recommended that all clones be rebuilt after upgrading.
Following code analysis caused by bug report by J Smith in message
CADFUPgc5bmtv-yg9znxV-vcfkb+JPRqs7m2OesQXaM_4Z1JpdQ@mail.gmail.com
and privately by F-Secure.
Backpatch to 9.3, where freezing of MultiXactIds was introduced.
Analysis and patch by Andres Freund, with some tweaks by Álvaro.
It is dangerous to do so, because some code expects to be able to see what's
the true Xmax even if it is aborted (particularly while traversing HOT
chains). So don't do it, and instead rely on the callers to verify for
abortedness, if necessary.
Several race conditions and bugs fixed in the process. One isolation test
changes the expected output due to these.
This also reverts commit c235a6a589, which is no longer necessary.
Backpatch to 9.3, where this function was introduced.
Andres Freund
Commit 9dc842f08 of 8.2 era prevented MultiXact truncation during crash
recovery, because there was no guarantee that enough state had been
setup, and because it wasn't deemed to be a good idea to remove data
during crash recovery anyway. Since then, due to Hot-Standby, streaming
replication and PITR, the amount of time a cluster can spend doing crash
recovery has increased significantly, to the point that a cluster may
even never come out of it. This has made not truncating the content of
pg_multixact/ not defensible anymore.
To fix, take care to setup enough state for multixact truncation before
crash recovery starts (easy since checkpoints contain the required
information), and move the current end-of-recovery actions to a new
TrimMultiXact() function, analogous to TrimCLOG().
At some later point, this should probably done similarly to the way
clog.c is doing it, which is to just WAL log truncations, but we can't
do that for the back branches.
Back-patch to 9.0. 8.4 also has the problem, but since there's no hot
standby there, it's much less pressing. In 9.2 and earlier, this patch
is simpler than in newer branches, because multixact access during
recovery isn't required. Add appropriate checks to make sure that's not
happening.
Andres Freund
While autovacuum dutifully launched anti-multixact-wraparound vacuums
when the multixact "age" was reached, the vacuum code was not aware that
it needed to make them be full table vacuums. As the resulting
partial-table vacuums aren't capable of actually increasing relminmxid,
autovacuum continued to launch anti-wraparound vacuums that didn't have
the intended effect, until age of relfrozenxid caused the vacuum to
finally be a full table one via vacuum_freeze_table_age.
To fix, introduce logic for multixacts similar to that for plain
TransactionIds, using the same GUCs.
Backpatch to 9.3, where permanent MultiXactIds were introduced.
Andres Freund, some cleanup by Álvaro
Parts of the code used autovacuum_freeze_max_age to determine whether
anti-multixact-wraparound vacuums are necessary, while others used a
hardcoded 200000000 value. This leads to problems when
autovacuum_freeze_max_age is set to a non-default value. Use the latter
everywhere.
Backpatch to 9.3, where vacuuming of multixacts was introduced.
Andres Freund
Prevent handle_sig_alarm from losing control partway through due to a query
cancel (either an asynchronous SIGINT, or a cancel triggered by one of the
timeout handler functions). That would at least result in failure to
schedule any required future interrupt, and might result in actual
corruption of timeout.c's data structures, if the interrupt happened while
we were updating those.
We could still lose control if an asynchronous SIGINT arrives just as the
function is entered. This wouldn't break any data structures, but it would
have the same effect as if the SIGALRM interrupt had been silently lost:
we'd not fire any currently-due handlers, nor schedule any new interrupt.
To forestall that scenario, forcibly reschedule any pending timer interrupt
during AbortTransaction and AbortSubTransaction. We can avoid any extra
kernel call in most cases by not doing that until we've allowed
LockErrorCleanup to kill the DEADLOCK_TIMEOUT and LOCK_TIMEOUT events.
Another hazard is that some platforms (at least Linux and *BSD) block a
signal before calling its handler and then unblock it on return. When we
longjmp out of the handler, the unblock doesn't happen, and the signal is
left blocked indefinitely. Again, we can fix that by forcibly unblocking
signals during AbortTransaction and AbortSubTransaction.
These latter two problems do not manifest when the longjmp reaches
postgres.c, because the error recovery code there kills all pending timeout
events anyway, and it uses sigsetjmp(..., 1) so that the appropriate signal
mask is restored. So errors thrown outside any transaction should be OK
already, and cleaning up in AbortTransaction and AbortSubTransaction should
be enough to fix these issues. (We're assuming that any code that catches
a query cancel error and doesn't re-throw it will do at least a
subtransaction abort to clean up; but that was pretty much required already
by other subsystems.)
Lastly, ProcSleep should not clear the LOCK_TIMEOUT indicator flag when
disabling that event: if a lock timeout interrupt happened after the lock
was granted, the ensuing query cancel is still going to happen at the next
CHECK_FOR_INTERRUPTS, and we want to report it as a lock timeout not a user
cancel.
Per reports from Dan Wood.
Back-patch to 9.3 where the new timeout handling infrastructure was
introduced. We may at some point decide to back-patch the signal
unblocking changes further, but I'll desist from that until we hear
actual field complaints about it.
Although user-defined relations can't be directly created in
pg_catalog, it's possible for them to end up there, because you can
create them in some other schema and then use ALTER TABLE .. SET SCHEMA
to move them there. Previously, such relations couldn't afterwards
be manipulated, because IsSystemRelation()/IsSystemClass() rejected
all attempts to modify objects in the pg_catalog schema, regardless
of their origin. With this patch, they now reject only those
objects in pg_catalog which were created at initdb-time, allowing
most operations on user-created tables in pg_catalog to proceed
normally.
This patch also adds new functions IsCatalogRelation() and
IsCatalogClass(), which is similar to IsSystemRelation() and
IsSystemClass() but with a slightly narrower definition: only TOAST
tables of system catalogs are included, rather than *all* TOAST tables.
This is currently used only for making decisions about when
invalidation messages need to be sent, but upcoming logical decoding
patches will find other uses for this information.
Andres Freund, with some modifications by me.
In the GIN incomplete-splits patch, I used BlockIdDatas to store the block
number of left and right children, when inserting a downlink after a split
to an internal page posting list page. But gin_desc thought they were stored
as BlockNumbers.
Instead of simply checking the KEYS_UPDATED bit, we need to check
whether each lock held on the future version of the tuple conflicts with
the lock we're trying to acquire.
Per bug report #8434 by Tomonari Katsumata
Not doing so causes us to traverse an update chain that has been broken
by concurrent page pruning. All other code that traverses update chains
uses this check as one of the cases in which to stop iterating, so
replicate it here too. Failure to do so leads to erroneous CLOG,
subtrans or multixact lookups.
Per discussion following the bug report by J Smith in
CADFUPgc5bmtv-yg9znxV-vcfkb+JPRqs7m2OesQXaM_4Z1JpdQ@mail.gmail.com
as diagnosed by Andres Freund.
If a transaction updates/deletes a tuple just before aborting, and a
concurrent transaction tries to prune the page concurrently, the pruner
may see HeapTupleSatisfiesVacuum return HEAPTUPLE_DELETE_IN_PROGRESS,
but a later call to HeapTupleGetUpdateXid() return InvalidXid. This
would cause an assertion failure in development builds, but would be
otherwise Mostly Harmless.
Fix by checking whether the updater Xid is valid before trying to apply
it as page prune point.
Reported by Andres in 20131124000203.GA4403@alap2.anarazel.de
The reason for the fetch failure is that the tuple was removed because
it was dead; so the failure is innocuous and can be ignored. Moreover,
there's no need for further work and we can return success to the caller
immediately. EvalPlanQualFetch is doing something very similar to this
already.
Report and test case from Andres Freund in
20131124000203.GA4403@alap2.anarazel.de
Replace it with an approach similar to what GiST uses: when a page is split,
the left sibling is marked with a flag indicating that the parent hasn't been
updated yet. When the parent is updated, the flag is cleared. If an insertion
steps on a page with the flag set, it will finish split before proceeding
with the insertion.
The post-recovery cleanup mechanism was never totally reliable, as insertion
to the parent could fail e.g because of running out of memory or disk space,
leaving the tree in an inconsistent state.
This also divides the responsibility of WAL-logging more clearly between
the generic ginbtree.c code, and the parts specific to entry and posting
trees. There is now a common WAL record format for insertions and deletions,
which is written by ginbtree.c, followed by tree-specific payload, which is
returned by the placetopage- and split- callbacks.
Separate the insertion payload from the more static portions of GinBtree.
GinBtree now only contains information related to searching the tree, and
the information of what to insert is passed separately.
Add root block number to GinBtree, instead of passing it around all the
functions as argument.
Split off ginFinishSplit() from ginInsertValue(). ginFinishSplit is
responsible for finding the parent and inserting the downlink to it.
Change SET LOCAL/CONSTRAINTS/TRANSACTION behavior outside of a
transaction block from error (post-9.3) to warning. (Was nothing in <=
9.3.) Also change ABORT outside of a transaction block from notice to
warning.
These bugs can cause data loss on standbys started with hot_standby=on at
the moment they start to accept read only queries, by marking committed
transactions as uncommited. The likelihood of such corruptions is small
unless the primary has a high transaction rate.
5a031a5556 fixed bugs in HS's startup logic
by maintaining less state until at least STANDBY_SNAPSHOT_PENDING state
was reached, missing the fact that both clog and subtrans are written to
before that. This only failed to fail in common cases because the usage
of ExtendCLOG in procarray.c was superflous since clog extensions are
actually WAL logged.
f44eedc3f0f347a856eea8590730769125964597/I then tried to fix the missing
extensions of pg_subtrans due to the former commit's changes - which are
not WAL logged - by performing the extensions when switching to a state
> STANDBY_INITIALIZED and not performing xid assignments before that -
again missing the fact that ExtendCLOG is unneccessary - but screwed up
twice: Once because latestObservedXid wasn't updated anymore in that
state due to the earlier commit and once by having an off-by-one error in
the loop performing extensions. This means that whenever a
CLOG_XACTS_PER_PAGE (32768 with default settings) boundary was crossed
between the start of the checkpoint recovery started from and the first
xl_running_xact record old transactions commit bits in pg_clog could be
overwritten if they started and committed in that window.
Fix this mess by not performing ExtendCLOG() in HS at all anymore since
it's unneeded and evidently dangerous and by performing subtrans
extensions even before reaching STANDBY_SNAPSHOT_PENDING.
Analysis and patch by Andres Freund. Reported by Christophe Pettus.
Backpatch down to 9.0, like the previous commit that caused this.
RecoveryIsInProgress() can be called very frequently. During normal
operation, it just checks a backend-local variable and returns quickly,
but during hot standby, it checks a spinlock-protected shared variable.
Those spinlock acquisitions can become a point of contention on a busy
hot standby system.
Replace the spinlock acquisition with a memory barrier.
Per discussion with Andres Freund, Ants Aasma and Merlin Moncure.
This patch adds the ability to write TABLE( function1(), function2(), ...)
as a single FROM-clause entry. The result is the concatenation of the
first row from each function, followed by the second row from each
function, etc; with NULLs inserted if any function produces fewer rows than
others. This is believed to be a much more useful behavior than what
Postgres currently does with multiple SRFs in a SELECT list.
This syntax also provides a reasonable way to combine use of column
definition lists with WITH ORDINALITY: put the column definition list
inside TABLE(), where it's clear that it doesn't control the ordinality
column as well.
Also implement SQL-compliant multiple-argument UNNEST(), by turning
UNNEST(a,b,c) into TABLE(unnest(a), unnest(b), unnest(c)).
The SQL standard specifies TABLE() with only a single function, not
multiple functions, and it seems to require an implicit UNNEST() which is
not what this patch does. There may be something wrong with that reading
of the spec, though, because if it's right then the spec's TABLE() is just
a pointless alternative spelling of UNNEST(). After further review of
that, we might choose to adopt a different syntax for what this patch does,
but in any case this functionality seems clearly worthwhile.
Andrew Gierth, reviewed by Zoltán Böszörményi and Heikki Linnakangas, and
significantly revised by me
Split off the portion of ginInsertValue that inserts the tuple to current
level into a separate function, ginPlaceToPage. ginInsertValue's charter
is now to recurse up the tree to insert the downlink, when a page split is
required.
This is in preparation for a patch to change the way incomplete splits are
handled, which will need to do these operations separately. And IMHO makes
the code more readable anyway.
This creates a new gin-btree callback function for creating a downlink for
a page. Previously, ginxlog.c duplicated the logic used during normal
operation.
The root page is filled with as many items as fit, and the rest are inserted
using normal insertions. However, I fumbled the variable names, and the code
actually memcpy'd all the items on the page, overflowing the buffer. While
at it, rename the variable to make the distinction more clear.
Reported by Teodor Sigaev. This bug was introduced by my recent
refactorings, so no backpatching required.
If a page is deleted, and reused for something else, just as a search is
following a rightlink to it from its left sibling, the search would continue
scanning whatever the new contents of the page are. That could lead to
incorrect query results, or even something more curious if the page is
reused for a different kind of a page.
To fix, modify the search algorithm to lock the next page before releasing
the previous one, and refrain from deleting pages from the leftmost branch
of the tree.
Add a new Concurrency section to the README, explaining why this works.
There is a lot more one could say about concurrency in GIN, but that's for
another patch.
Backpatch to all supported versions.
Merge the isEnoughSpace and placeToPage functions in the b-tree interface
into one function that tries to put a tuple on page, and returns false if
it doesn't fit.
Move createPostingTree function to gindatapage.c, and change its contract
so that it can be passed more items than fit on the root page. It's in a
better position than the callers to know how many items fit.
Move ginMergeItemPointers out of gindatapage.c, into a separate file.
These changes make no difference now, but reduce the footprint of Alexander
Korotkov's upcoming patch to pack item pointers more tightly.
Historically, printtup() has assumed that it could prevent memory leakage
by pfree'ing the string result of each output function and manually
managing detoasting of toasted values. This amounts to assuming that
datatype output functions never leak any memory internally; an assumption
we've already decided to be bogus elsewhere, for example in COPY OUT.
range_out in particular is known to leak multiple kilobytes per call, as
noted in bug #8573 from Godfried Vanluffelen. While we could go in and fix
that leak, it wouldn't be very notationally convenient, and in any case
there have been and undoubtedly will again be other leaks in other output
functions. So what seems like the best solution is to run the output
functions in a temporary memory context that can be reset after each row,
as we're doing in COPY OUT. Some quick experimentation suggests this is
actually a tad faster than the retail pfree's anyway.
This patch fixes all the variants of printtup, except for debugtup()
which is used in standalone mode. It doesn't seem worth worrying
about query-lifespan leaks in standalone mode, and fixing that case
would be a bit tedious since debugtup() doesn't currently have any
startup or shutdown functions.
While at it, remove manual detoast management from several other
output-function call sites that had copied it from printtup(). This
doesn't make a lot of difference right now, but in view of recent
discussions about supporting "non-flattened" Datums, we're going to
want that code gone eventually anyway.
Back-patch to 9.2 where range_out was introduced. We might eventually
decide to back-patch this further, but in the absence of known major
leaks in older output functions, I'll refrain for now.
The original coding thought this case was impossible, but it can happen
if the bgwriter or checkpointer processes decide to write out an index
page while creation is still proceeding, leading to a bogus "unexpected
spgdoinsert() failure" error. Problem reported by Jonathan S. Katz.
Teodor Sigaev
The C and POSIX standards state that strncpy's behavior is undefined when
source and destination areas overlap. While it remains dubious whether any
implementations really misbehave when the pointers are exactly equal, some
platforms are now starting to force the issue by complaining when an
undefined call occurs. (In particular OS X 10.9 has been seen to dump core
here, though the exact set of circumstances needed to trigger that remain
elusive. Similar behavior can be expected to be optional on Linux and
other platforms in the near future.) So tweak the code to explicitly do
nothing when nothing need be done.
Back-patch to all active branches. In HEAD, this also lets us get rid of
an exception in valgrind.supp.
Per discussion of a report from Matthias Schmitt.
This avoids an assumption about the signed number representation. It is
anticipated to have no functional changes on supported configurations;
many two's complement assumptions remain elsewhere.
Per a suggestion from Andres Freund.
The TYPEALIGN macro, and the related ones like MAXALIGN, don't work with
values larger than intptr_t, because TYPEALIGN casts the argument to
intptr_t to do the arithmetic. That's not a problem when dealing with
pointers or lengths or offsets related to pointers, but the XLogInsert
scaling patch added a call to MAXALIGN with an XLogRecPtr argument.
To fix, add wider variants of the macros, called TYPEALIGN64 and MAXALIGN64,
which are just like the existing variants but work with uint64 instead of
intptr_t.
Report and patch by David Rowley, analysis by Andres Freund.
1. In heap_hot_search_buffer(), the PredicateLockTuple() call is passed
wrong offset number. heapTuple->t_self is set to the tid of the first
tuple in the chain that's visited, not the one actually being read.
2. CheckForSerializableConflictIn() uses the tuple's t_ctid field
instead of t_self to check for exiting predicate locks on the tuple. If
the tuple was updated, but the updater rolled back, t_ctid points to the
aborted dead tuple.
Reported by Hannu Krosing. Backpatch to 9.1.
It makes for cleaner code to have separate Get/Add functions for PostingItems
and ItemPointers. A few callsites that have to deal with both types need to
be duplicated because of this, but all the callers have to know which one
they're dealing with anyway. Overall, this reduces the amount of casting
required.
Extracted from Alexander Korotkov's larger patch to change the data page
format.
It seems to make more sense to use "cutoff multixact" terminology
throughout the backend code; "freeze" is associated with replacing of an
Xid with FrozenTransactionId, which is not what we do for MultiXactIds.
Andres Freund
Some adjustments by Álvaro Herrera
This reverts commit 269e780822
and commit 5b571bb8c8.
Unfortunately, the initial patch had insufficient performance testing,
and resulted in a regression.
Per report by Thom Brown.
Performance testing shows that if the insertpos_lck spinlock and the fields
that it protects are on the same cache line with other variables that are
frequently accessed, the false sharing can hurt performance a lot. Keep
them apart by adding some padding.
This keeps the usual trigger file name unchanged from 9.2, avoiding nasty
issues if you use a pre-9.3 pg_ctl binary with a 9.3 server or vice versa.
The fallback behavior of creating a full checkpoint before starting up is now
triggered by a file called "fallback_promote". That can be useful for
debugging purposes, but we don't expect any users to have to resort to that
and we might want to remove that in the future, which is why the fallback
mechanism is undocumented.
When upgrading from servers of versions 9.2 and older, and MultiXactIds
have been used in the old server beyond the first page (that is, 2048
multis or more in the default 8kB-page build), pg_upgrade would set the
next multixact offset to use beyond what has been allocated in the new
cluster. This would cause a failure the first time the new cluster
needs to use this value, because the pg_multixact/offsets/ file wouldn't
exist or wouldn't be large enough. To fix, ensure that the transient
server instances launched by pg_upgrade extend the file as necessary.
Per report from Jesse Denardo in
CANiVXAj4c88YqipsyFQPboqMudnjcNTdB3pqe8ReXqAFQ=HXyA@mail.gmail.com
As far as I can determine, there's no code in the core distribution
that fails to explicitly set the snapshot of a scan or executor
state. If there is any such code, this will probably cause it to
seg fault; friendlier suggestions were discussed on pgsql-hackers,
but there was no consensus that anything more than this was
needed.
This is another step towards the hoped-for complete removal of
SnapshotNow.
Previously, these functions took a HeapTupleHeader, but upcoming
patches for logical replication will introduce new a new snapshot
type under which the tuple's TID will be used to lookup (CMIN, CMAX)
for visibility determination purposes. This makes that information
available. Code churn is minimal since HeapTupleSatisfiesVisibility
took the HeapTuple anyway, and deferenced it before calling the
satisfies function.
Independently of logical replication, this allows t_tableOid and
t_self to be cross-checked via assertions in tqual.c. This seems
like a useful way to make sure that all callers are setting these
values properly, which has been previously put forward as
desirable.
Andres Freund, reviewed by Álvaro Herrera
For simple views which are automatically updatable, this patch allows
the user to specify what level of checking should be done on records
being inserted or updated. For 'LOCAL CHECK', new tuples are validated
against the conditionals of the view they are being inserted into, while
for 'CASCADED CHECK' the new tuples are validated against the
conditionals for all views involved (from the top down).
This option is part of the SQL specification.
Dean Rasheed, reviewed by Pavel Stehule
Initialization of the first XLOG buffer at end-of-recovery was broken for
the case that the last read WAL record ended at a page boundary. Instead of
trying to copy the last full xlog page to the buffer cache in that case,
just set shared state so that the next page is initialized when the first
WAL record after startup is inserted. (that's what we did in earlier
version, too)
To make the shared state required for that case less surprising, replace the
XLogCtl->curridx variable, which was the index of the latest initialized
buffer, with an XLogRecPtr of how far the buffers have been initialized.
That also allows us to get rid of the XLogRecEndPtrToBufIdx macro.
While we're at it, make a similar change for XLogCtl->Write.curridx, getting
rid of that variable and calculating the next buffer to write from
XLogCtl->LogwrtResult instead.
Since this function assumed non-MVCC snapshots, it broke when commit
568d4138c6 switched its one caller from
SnapshotNow scans to MVCC-snapshot scans.
Reviewed by Robert Haas, Tom Lane and Andres Freund.
Was broken by my xloginsert scaling patch. XLogCtl global variable needs
to be initialized in each process, as it's not inherited by fork() on
Windows.
This patch replaces WALInsertLock with a number of WAL insertion slots,
allowing multiple backends to insert WAL records to the WAL buffers
concurrently. This is particularly useful for parallel loading large amounts
of data on a system with many CPUs.
This has one user-visible change: switching to a new WAL segment with
pg_switch_xlog() now fills the remaining unused portion of the segment with
zeros. This potentially adds some overhead, but it has been a very common
practice by DBA's to clear the "tail" of the segment with an external
pg_clearxlogtail utility anyway, to make the WAL files compress better.
With this patch, it's no longer necessary to do that.
This patch adds a new GUC, xloginsert_slots, to tune the number of WAL
insertion slots. Performance testing suggests that the default, 8, works
pretty well for all kinds of worklods, but I left the GUC in place to allow
others with different hardware to test that easily. We might want to remove
that before release.
Reviewed by Andres Freund.
On some platforms, posix_fallocate() is available but may still return
EINVAL if the underlying filesystem does not support it. So, in case
of an error, fall through to the alternate implementation that just
writes zeros.
Per buildfarm failure and analysis by Tom Lane.
This function is more efficient than actually writing out zeroes to
the new file, per microbenchmarks by Jon Nelson. Also, it may reduce
the likelihood of WAL file fragmentation.
Jon Nelson, with review by Andres Freund, Greg Smith and me.
In 9.3, there's no particular limit on the number of bgworkers;
instead, we just count up the number that are actually registered,
and use that to set MaxBackends. However, that approach causes
problems for Hot Standby, which needs both MaxBackends and the
size of the lock table to be the same on the standby as on the
master, yet it may not be desirable to run the same bgworkers in
both places. 9.3 handles that by failing to notice the problem,
which will probably work fine in nearly all cases anyway, but is
not theoretically sound.
A further problem with simply counting the number of registered
workers is that new workers can't be registered without a
postmaster restart. This is inconvenient for administrators,
since bouncing the postmaster causes an interruption of service.
Moreover, there are a number of applications for background
processes where, by necessity, the background process must be
started on the fly (e.g. parallel query). While this patch
doesn't actually make it possible to register new background
workers after startup time, it's a necessary prerequisite.
Patch by me. Review by Michael Paquier.
Treat TOAST index just the same as normal one and get the OID
of TOAST index from pg_index but not pg_class.reltoastidxid.
This change allows us to handle multiple TOAST indexes, and
which is required infrastructure for upcoming
REINDEX CONCURRENTLY feature.
Patch by Michael Paquier, reviewed by Andres Freund and me.
To that end, support tags rather than lengths for external datums.
As an example of how this can be used, add support or "indirect"
tuples which point to some externally allocated memory containing
a toast tuple. Similar infrastructure could be used for other
purposes, including, perhaps, support for alternative compression
algorithms.
Andres Freund, reviewed by Hitoshi Harada and myself
SnapshotNow scans have the undesirable property that, in the face of
concurrent updates, the scan can fail to see either the old or the new
versions of the row. In many cases, we work around this by requiring
DDL operations to hold AccessExclusiveLock on the object being
modified; in some cases, the existing locking is inadequate and random
failures occur as a result. This commit doesn't change anything
related to locking, but will hopefully pave the way to allowing lock
strength reductions in the future.
The major issue has held us back from making this change in the past
is that taking an MVCC snapshot is significantly more expensive than
using a static special snapshot such as SnapshotNow. However, testing
of various worst-case scenarios reveals that this problem is not
severe except under fairly extreme workloads. To mitigate those
problems, we avoid retaking the MVCC snapshot for each new scan;
instead, we take a new snapshot only when invalidation messages have
been processed. The catcache machinery already requires that
invalidation messages be sent before releasing the related heavyweight
lock; else other backends might rely on locally-cached data rather
than scanning the catalog at all. Thus, making snapshot reuse
dependent on the same guarantees shouldn't break anything that wasn't
already subtly broken.
Patch by me. Review by Michael Paquier and Andres Freund.
We don't normally bother retrying when the number of bytes written by
write() is short of what was requested. It is generally assumed that a
write() to disk doesn't return short, unless you run out of disk space.
While writing the WAL, however, it seems prudent to try a bit harder,
because a failure leads to PANIC. The write() is also much larger than most
write()s in the backend (up to wal_buffers), so there's more room for
surprises.
Also retry on EINTR. All signals used in the backend are flagged SA_RESTART
nowadays, so it shouldn't happen, but better to be defensive.
Valgrind "client requests" in aset.c and mcxt.c teach Valgrind and its
Memcheck tool about the PostgreSQL allocator. This makes Valgrind
roughly as sensitive to memory errors involving palloc chunks as it is
to memory errors involving malloc chunks. Further client requests in
PageAddItem() and printtup() verify that all bits being added to a
buffer page or furnished to an output function are predictably-defined.
Those tests catch failures of C-language functions to fully initialize
the bits of a Datum, which in turn stymie optimizations that rely on
_equalConst(). Define the USE_VALGRIND symbol in pg_config_manual.h to
enable these additions. An included "suppression file" silences nominal
errors we don't plan to fix.
Reviewed in earlier versions by Peter Geoghegan and Korry Douglas.
Every other core buffer page consumer initializes the bytes it furnishes
to PageAddItem(). For consistency, do the same here. No back-patch;
regardless, we couldn't count on the fix so long as binary upgrade can
carry forward affected index builds.
Clang 3.3 correctly complains that a variable of type enum
MultiXactStatus cannot hold a value of -1, which makes sense. Change
the declared type of the variable to int instead, and apply casting as
necessary to avoid the warning.
Per notice from Andres Freund
In some cases with higher numbers of subtransactions
it was possible for us to incorrectly initialize
subtrans leading to complaints of missing pages.
Bug report by Sergey Konoplev
Analysis and fix by Andres Freund
Most of the documentation uses "single-user mode", so use that in the
code as well. Adjust the documentation to match the new error message
wording. Also add a documentation index entry for "single-user mode".
Based-on-patch-by: Jeff Janes <jeff.janes@gmail.com>
MarkBufferDirtyHint() writes WAL, and should know if it's got a
standard buffer or not. Currently, the only callers where buffer_std
is false are related to the FSM.
In passing, rename XLOG_HINT to XLOG_FPI, which is more descriptive.
Back-patch to 9.3.
SP-GiST's original scheme for avoiding deadlocks during concurrent index
insertions doesn't work, as per report from Hailong Li, and there isn't any
evident way to make it work completely. We could possibly lock individual
inner tuples instead of their whole pages, but preliminary experimentation
suggests that the performance penalty would be huge. Instead, if we fail
to get a buffer lock while descending the tree, just restart the tree
descent altogether. We keep the old tuple positioning rules, though, in
hopes of reducing the number of cases where this can happen.
Teodor Sigaev, somewhat edited by Tom Lane
elog.c has historically treated LOG messages as low-priority during
bootstrap and standalone operation. This has led to confusion and even
masked a bug, because the normal expectation of code authors is that
elog(LOG) will put something into the postmaster log, and that wasn't
happening during initdb. So get rid of the special-case rule and make
the priority order the same as it is in normal operation. To keep from
cluttering initdb's output and the behavior of a standalone backend,
tweak the severity level of three messages routinely issued by xlog.c
during startup and shutdown so that they won't appear in these cases.
Per my proposal back in December.
Since commit f21bb9cfb5, this function
ignores the caller-provided length and loops until it finds a
terminator, which GetVirtualXIDsDelayingChkpt() never adds. Restore the
previous loop control logic. In passing, revert the addition of an
unused variable by the same commit, presumably a debugging relic.
Seems cleaner to get the currently-replayed TLI in the same call to
GetXLogReplayRecPtr that we get the WAL position. Make it more clear in the
comment what the code does when recovery has already ended
(RecoveryInProgress() will set ThisTimeLineID in that case). Finally, make
resetting ThisTimeLineID afterwards more explicit.
Make slightly better decisions about indentation than what pgindent
is capable of. Mostly breaking out long function calls into one
line per argument, with a few other minor adjustments.
No functional changes- all whitespace.
pgindent ran cleanly (didn't change anything) after.
Passes all regressions.
checkpointer needs to reset ThisTimeLineID after
a restartpoint to allow installing/recycling new
WAL files. If recovery has already ended this
would leave ThisTimeLineID set incorrectly and
so we must reset it otherwise later checkpoints
do not have the correct timeline.
Bug report by Heikki Linnakangas.
Further investigation by Heikki and myself.
This simplifies the handling of crashes after fast promotion and various
minor cases that can exist in short timing windows around that case.
Broad fix to bug reported by Michael Paquier on -hackers,
approach prompted by Heikki Linnakangas
If OID wraparound should occur while in standalone mode (unlikely but
possible), we want to advance the counter to FirstNormalObjectId not
FirstBootstrapObjectId. Otherwise, user objects might be created with OIDs
in the system-reserved range. That isn't immediately harmful but it poses
a risk of conflicts during future pg_upgrade operations.
Noted by Andres Freund. Back-patch to all supported branches, since all of
them are supported sources for pg_upgrade operations.
Commit d22a09dc70 introduced official support
for GiST consistentFns that want to cache data using the FmgrInfo fn_extra
pointer: the idea was to preserve the cached values across gistrescan(),
whereas formerly they'd been leaked. However, there was an oversight in
that, namely that multiple scan keys might reference the same column's
consistentFn; the code would result in propagating the same cache value
into multiple scan keys, resulting in crashes or wrong answers. Use a
separate array instead to ensure that each scan key keeps its own state.
Per bug #8143 from Joel Roller. Back-patch to 9.2 where the bug was
introduced.
If a standby server has a cascading standby server connected to it, it's
possible that WAL has already been sent up to the next WAL page boundary,
splitting a WAL record in the middle, when the first standby server is
promoted. Don't throw an assertion failure or error in walsender if that
happens.
Also, fix a variant of the same bug in pg_receivexlog: if it had already
received WAL on previous timeline up to a segment boundary, when the
upstream standby server is promoted so that the timeline switch record falls
on the previous segment, pg_receivexlog would miss the segment containing
the timeline switch. To fix that, have walsender send the position of the
timeline switch at end-of-streaming, in addition to the next timeline's ID.
It was previously assumed that the switch happened exactly where the
streaming stopped.
Note: this is an incompatible change in the streaming protocol. You might
get an error if you try to stream over timeline switches, if the client is
running 9.3beta1 and the server is more recent. It should be fine after a
reconnect, however.
Reported by Fujii Masao.
What we have implemented is a radix tree (or a radix trie or a patricia
trie), but the docs and code comments incorrectly called it a "suffix tree".
Alexander Korotkov
This saves some memory from each index relcache entry. At least on a 64-bit
machine, it saves just enough to shrink a typical relcache entry's memory
usage from 2k to 1k. That's nice if you have a lot of backends and a lot of
indexes.
Heikki reported comment was wrong, so fixed
code to match the comment: we only need to
take additional locking precautions when we
have a shared lock on the buffer.
We copy the buffer before inserting an XLOG_HINT to avoid WAL CRC errors
caused by concurrent hint writes to buffer while share locked. To make this work
we refactor RestoreBackupBlock() to allow an XLOG_HINT to avoid the normal
path for backup blocks, which assumes the underlying buffer is exclusive locked.
Resulting code completely changes layout of XLOG_HINT WAL records, but
this isn't even beta code, so this is a low impact change.
In passing, avoid taking WALInsertLock for full page writes on checksummed
hints, remove related cruft from XLogInsert() and improve xlog_desc record for
XLOG_HINT.
Andres Freund
Bug report by Fujii Masao, testing by Jeff Janes and Jaime Casanova,
review by Jeff Davis and Simon Riggs. Applied with changes from review
and some comment editing.
In CLUSTER, VACUUM FULL and ALTER TABLE SET TABLESPACE
I erroneously set checksum before log_newpage, which
sets the LSN and invalidates the checksum. So set
checksum immediately *after* log_newpage.
Bug report Fujii Masao, Fix and patch by Jeff Davis
The pg_start_backup() and pg_stop_backup() functions checked the privileges
of the initially-authenticated user rather than the current user, which is
wrong. For example, a user-defined index function could successfully call
these functions when executed by ANALYZE within autovacuum. This could
allow an attacker with valid but low-privilege database access to interfere
with creation of routine backups. Reported and fixed by Noah Misch.
Security: CVE-2013-1901
In a heap update, if the old and new tuple were on different pages, and the
new page no longer existed (because it was subsequently truncated away by
vacuum), heap_xlog_update forgot to release the pin on the old buffer. This
bug was introduced by the "Fix multiple problems in WAL replay" patch,
commit 3bbf668de9 (on master branch).
With full_page_writes=off, this triggered an "incorrect local pin count"
error later in replay, if the old page was vacuumed.
This fixes bug #7969, reported by Yunong Xiao. Backpatch to 9.0, like the
commit that introduced this bug.
Checksums are set immediately prior to flush out of shared buffers
and checked when pages are read in again. Hint bit setting will
require full page write when block is dirtied, which causes various
infrastructure changes. Extensive comments, docs and README.
WARNING message thrown if checksum fails on non-all zeroes page;
ERROR thrown but can be disabled with ignore_checksum_failure = on.
Feature enabled by an initdb option, since transition from option off
to option on is long and complex and has not yet been implemented.
Default is not to use checksums.
Checksum used is WAL CRC-32 truncated to 16-bits.
Simon Riggs, Jeff Davis, Greg Smith
Wide input and assistance from many community members. Thank you.
Remove use of PageSetTLI() from all page manipulation functions
and adjust README to indicate change in the way we make changes
to pages. Repurpose those bytes into the pd_checksum field and
explain how that works in comments about page header.
Refactoring ahead of actual feature patch which would make use
of the checksum field, arriving later.
Jeff Davis, with comments and doc changes by Simon Riggs
Direction suggested by Robert Haas; many others providing
review comments.
We had two copies of this function in the backend and libpq, which was
already pretty bogus, but it turns out that we need it in some other
programs that don't use libpq (such as pg_test_fsync). So put it where
it probably should have been all along. The signal-mask-initialization
support in src/backend/libpq/pqsignal.c stays where it is, though, since
we only need that in the backend.
If we were about to enter archive recovery after crash recovery, we scanned
the archive for the latest tli history file, and set the recovery target
timeline to that. However, when we actually tried to read the history file,
we would not fetch the file from the archive, because we were not in archive
recovery yet.
To fix, make readTimeLineHistory and existsTimeLineHistory to always fetch
the file from archive if archive recovery is requested, even if we're not in
archive recovery yet.
Backpatch to 9.2. Mitsumasa KONDO
A materialized view has a rule just like a view and a heap and
other physical properties like a table. The rule is only used to
populate the table, references in queries refer to the
materialized data.
This is a minimal implementation, but should still be useful in
many cases. Currently data is only populated "on demand" by the
CREATE MATERIALIZED VIEW and REFRESH MATERIALIZED VIEW statements.
It is expected that future releases will add incremental updates
with various timings, and that a more refined concept of defining
what is "fresh" data will be developed. At some point it may even
be possible to have queries use a materialized in place of
references to underlying tables, but that requires the other
above-mentioned features to be working first.
Much of the documentation work by Robert Haas.
Review by Noah Misch, Thom Brown, Robert Haas, Marko Tiikkaja
Security review by KaiGai Kohei, with a decision on how best to
implement sepgsql still pending.
fmgr_sql had been designed on the assumption that the FmgrInfo it's called
with has only query lifespan. This is demonstrably unsafe in connection
with range types, as shown in bug #7881 from Andrew Gierth. Fix things
so that we re-generate the function's cache data if the (sub)transaction
it was made in is no longer active.
Back-patch to 9.2. This might be needed further back, but it's not clear
whether the case can realistically arise without range types, so for now
I'll desist from back-patching further.
We must still initialize minRecoveryPoint if we start straight with archive
recovery, e.g when recovering from a normal base backup taken with
pg_start/stop_backup. Otherwise we never consider the system consistent.
If you create a base backup using an atomic filesystem snapshot, and try to
perform PITR starting from that base backup, or if you just kill a master
server and create recovery.conf to put it into standby mode, we don't know
how far we need to recover before reaching consistency. Normally in crash
recovery, we replay all the WAL present in pg_xlog, and assume that we're
consistent after that. And normally in archive recovery, minRecoveryPoint,
backupEndRequired, or backupEndPoint is set in the control file, indicating
how far we need to replay to reach consistency. But if the server was
previously up and running normally, and you kill -9 it or take an atomic
filesystem snapshot, none of those fields are set in the control file.
The solution is to perform crash recovery first, replaying all the WAL in
pg_xlog. After that's done, we assume that the system is consistent like in
normal crash recovery, and switch to archive recovery mode after that.
Per report from Kyotaro HORIGUCHI. In his scenario, recovery.conf was
created after "pg_ctl stop -m i". I'm not sure we need to support that exact
scenario, but we should support backing up using a filesystem snapshot,
which looks identical.
This issue goes back to at least 9.0, where hot standby was introduced and
we started to track when consistency is reached. In 9.1 and 9.2, we would
open up for hot standby too early, and queries could briefly see an
inconsistent state. But 9.2 made it more visible, as we started to PANIC if
we see a reference to a non-existing page during recovery, if we've already
reached consistency. This is a fairly big patch, so back-patch to 9.2 only,
where the issue is more visible. We can consider back-patching further after
this has received some more testing in 9.2 and master.
This enables non-backend code, such as pg_xlogdump, to use it easily.
The previous location, in src/backend/catalog/catalog.c, made that
essentially impossible because that file depends on many backend-only
facilities; so this needs to live separately.
Revert my earlier fix for the bug that unarchived WAL files get deleted on
crash recovery, commit c9cc7e05c6. We create
a .done file for files streamed or restored from archive, so the WAL file
recycling logic used during normal operation works just as well during
archive recovery.
Per Fujii Masao's suggestion.
This is a forward-patch of commit 6f4b8a4f4f,
applied to 9.2 back in August. The plan was to do something else in master,
but it looks like it's not going to happen, so let's just apply the 9.2
solution to master as well.
Fujii Masao
Currently it's only possible for loadable modules to get control during
post-commit cleanup of a transaction. That doesn't work too well if they
want to do something that could throw an error; for example, an FDW might
need to issue a remote commit, which could well fail. To improve matters,
extend the existing APIs for XactCallback and SubXactCallback functions
to provide new pre-commit events for this purpose.
The release notes will need to mention that existing callback functions
should be checked to make sure they don't do something unwanted when one
of the new event types occurs. In the examples within our source tree,
contrib/sepgsql was fine but plpgsql had been a bit too cute.
The reason this wasn't supported before was that GiST indexes need an
increasing sequence to detect concurrent page-splits. In a regular WAL-
logged GiST index, the LSN of the page-split record is used for that
purpose, and in a temporary index, we can get away with a backend-local
counter. Neither of those methods works for an unlogged relation.
To provide such an increasing sequence of numbers, create a "fake LSN"
counter that is saved and restored across shutdowns. On recovery, unlogged
relations are blown away, so the counter doesn't need to survive that
either.
Jeevan Chalke, based on discussions with Robert Haas, Tom Lane and me.
The intention was to request a regular online checkpoint immediately after
end of recovery, when performing "fast promotion". However, because the
checkpoint was requested before other backends were allowed to write WAL,
the checkpointer process performed a restartpoint rather than a checkpoint.
Delay the RequestCheckPoint call until after recovery has truly ended, so
that you get a real checkpoint.
This isn't used for anything but a sanity check at the moment, but it could
be highly valuable for debugging purposes. It could also be used to recreate
timeline history by traversing WAL, which seems useful.
After further reflection I was unconvinced that the existing coding is
guaranteed to return valid union datums in every code path for multi-column
indexes. Fix that by forcing a gistunionsubkey() call at the end of the
recursion. Having done that, we can remove some clearly-redundant calls
elsewhere. This should be a little faster for multi-column indexes (since
the previous coding would uselessly do such a call for each column while
unwinding the recursion), as well as much harder to break.
Also, simplify the handling of cases where one side or the other of a
primary split contains only don't-care tuples. The previous coding used a
very ugly hack in removeDontCares() that essentially forced one random
tuple to be treated as non-don't-care, providing a random initial choice of
seed datum for the secondary split. It seems unlikely that that method
will give better-than-random splits. Instead, treat such a split as
degenerate and just let the next column determine the split, the same way
that we handle fully degenerate cases where the two sides produce identical
union datums.
This LOG message was put in over five years ago with the evident
expectation that we'd make all GiST opclasses support secondary split
directly. However, no such thing ever happened, and indeed the number of
opclasses supporting it decreased to zero in 9.2. The reason is that
improving on the default implementation isn't that easy --- the
opclass-specific code that did exist, before 9.2, doesn't appear to have
been any improvement over the default.
Hence, remove the message altogether. There's certainly no point in
nagging users about this in released branches, but I doubt that we'll
ever implement complete opclass-specific support anyway.
Not only is this implementation of secondary-split not better than the
default implementation in gistsplit.c, it's actually worse. The gistsplit.c
code at least looks to see if switching the left and right sides would make
a better merge with the previously-split tuples, while this doesn't.
In any case it's rather useless to support secondary split only in an edge
case. There used to be more complete support for it here (in chooseLR()),
but that was removed in commit 7f3bd86843.
It appears to me though that the chooseLR() code was really isomorphic to
the default implementation, since it was still based on choosing the cheaper
way of adding two sub-split vectors that had been chosen without regard to
the primary split initially. I think an implementation of secondary split
that could beat the default implementation would have to be pretty fully
integrated into the split algorithm, not plastered on at the end.
Back-patch to 9.2, but not further; previous branches have the chooseLR()
code which I don't feel a great need to mess with. This is mainly so we
just have two behaviors and not three among the various branches (IOW, this
patch is cleanup for commit 7f3bd86843e5aad84585a57d3f6b80db3c609916's
incomplete removal of secondary-split support).
Improve comments, rename some variables and functions, slightly simplify
a couple of APIs, in an attempt to make this code readable by people other
than its original author.
Even though this is essentially just cosmetic, back-patch to all active
branches, because otherwise it's going to make back-patching future fixes
in this file very painful.
While there's considerable doubt that we want fuzzy behavior in the
geometric operators at all (let alone as currently implemented), nobody is
stepping forward to redesign that stuff. In the meantime it behooves us
to make sure that index searches agree with the behavior of the underlying
operators. This patch fixes two problems in this area.
First, gist_box_same was using fuzzy equality, but it really needs to use
exact equality to prevent not-quite-identical upper index keys from being
treated as identical, which for example would prevent an existing upper
key from being extended by an amount less than epsilon. This would result
in inconsistent indexes. (The next release notes will need to recommend
that users reindex GiST indexes on boxes, polygons, circles, and points,
since all four opclasses use gist_box_same.)
Second, gist_point_consistent used exact comparisons for upper-page
comparisons in ~= searches, when it needs to use fuzzy comparisons to
ensure it finds all matches; and it used fuzzy comparisons for point <@ box
searches, when it needs to use exact comparisons because that's what the
<@ operator (rather inconsistently) does.
The added regression test cases illustrate all three misbehaviors.
Back-patch to all active branches. (8.4 did not have GiST point_ops,
but it still seems prudent to apply the gist_box_same patch to it.)
Alexander Korotkov, reviewed by Noah Misch
When considering a non-last column in a multi-column GiST index,
gistsplit.c tries to improve on the split chosen by the opclass-specific
pickSplit function by considering penalties for the next column. However,
there were two bugs in this code: it failed to recompute the union keys for
the leftmost index columns, even though these might well change after
reassigning tuples; and it included the old union keys in the recomputation
for the columns it did recompute, so that those keys couldn't get smaller
even if they should. The first problem could result in an invalid index
in which searches wouldn't find index entries that are in fact present;
the second would make the index less efficient to search.
Both of these errors were caused by misuse of gistMakeUnionItVec, whose
API was designed in a way that just begged such errors to be made. There
is no situation in which it's safe or useful to compute the union keys for
a subset of the index columns, and there is no caller that wants any
previous union keys to be included in the computation; so the undocumented
choice to treat the union keys as in/out rather than pure output parameters
is a waste of code as well as being dangerous.
Hence, rather than just making a minimal patch, I've changed the API of
gistMakeUnionItVec to remove the "startkey" parameter (it now always
processes all index columns) and treat the attr/isnull arrays as purely
output parameters.
In passing, also get rid of a couple of unnecessary and dangerous uses
of static variables in gistutil.c. It's remarkable that the one in
gistMakeUnionKey hasn't given us portability troubles before now, because
in addition to posing a re-entrancy hazard, it was unsafely assuming that
a static char[] array would have at least Datum alignment.
Per investigation of a trouble report from Tomas Vondra. (There are also
some bugs in contrib/btree_gist to be fixed, but that seems like material
for a separate patch.) Back-patch to all supported branches.
The new rmgrlist.h header, containing all necessary data
about built-in resource managers, allows other pieces of code to
access them.
In particular, this allows a future pg_xlogdump program to extract
rm_desc function pointers, without having to keep a duplicate list of
them.
Failing to do this results in almost all updates to system catalogs
being non-HOT updates, because the OID column would differ (not having
been set for the new tuple), which is an indexed column.
While at it, make sure to set the tableoid early in both old and new
tuples as well. This isn't of much consequence, since that column is
seldom (never?) indexed.
Report and patch from Andres Freund.
We must only set the bit(s) for the strongest lock held in the tuple;
otherwise, a multixact containing members with exclusive lock and
key-share lock will behave as though only a share lock is held.
This bug was introduced in commit 0ac5ad5134, somewhere along
development, when we allowed a singleton FOR SHARE lock to be
implemented without a MultiXact by using a multi-bit pattern.
I overlooked that GetMultiXactIdHintBits() needed to be tweaked as well.
Previously, we could have the bits for FOR KEY SHARE and FOR UPDATE
simultaneously set and it wouldn't cause a problem.
Per report from digoal@126.com
Previous patch to skip checkpoints at end of recovery didn't
correctly perform crash recovery, fumbling the timeline switch.
Now we record the minRecoveryPointTLI of the newly selected
timeline, so that we crash recover to the correct timeline.
Bug report from Fujii Masao, investigated by me.
This patch addresses the problem that applications currently have to
extract object names from possibly-localized textual error messages,
if they want to know for example which index caused a UNIQUE_VIOLATION
failure. It adds new error message fields to the wire protocol, which
can carry the name of a table, table column, data type, or constraint
associated with the error. (Since the protocol spec has always instructed
clients to ignore unrecognized field types, this should not create any
compatibility problem.)
Support for providing these new fields has been added to just a limited set
of error reports (mainly, those in the "integrity constraint violation"
SQLSTATE class), but we will doubtless add them to more calls in future.
Pavel Stehule, reviewed and extensively revised by Peter Geoghegan, with
additional hacking by Tom Lane.
pg_ctl promote -m fast will skip the checkpoint at end of recovery so that we
can achieve very fast failover when the apply delay is low. Write new WAL record
XLOG_END_OF_RECOVERY to allow us to switch timeline correctly for downstream log
readers. If we skip synchronous end of recovery checkpoint we request a normal
spread checkpoint so that the window of re-recovery is low.
Simon Riggs and Kyotaro Horiguchi, with input from Fujii Masao.
Review by Heikki Linnakangas
When descending the tree for an insert, and there are multiple equally good
pages we could insert to, make the choice in random. Previously, we would
always choose the tuple with lowest offset number. That meant that when two
non-leaf pages overlap - in the extreme case they might have exactly the same
key - all but the first such page went unused. That wasn't optimal for space
usage; if you deleted some tuples from the non-first pages, the space would
never be reused.
With this patch, the other pages are sometimes chosen too, although there's
still a heavy bias towards low-offset tuples, so that we don't lose cache
locality when doing a lot of inserts with similar keys.
Original idea by Alexander Korotkov, although this patch version was written
by me and copy-edited by Tom Lane.
The machinery around XLOG_HEAP2_CLEANUP_INFO failed
to correctly pass through the necessary information
on latestRemovedXid, avoiding cancellations in some
infrequent concurrent update/cleanup scenarios.
Backpatchable fix to 9.0
Detailed bug report and fix by Noah Misch,
backpatchable version by me.
This patch introduces two additional lock modes for tuples: "SELECT FOR
KEY SHARE" and "SELECT FOR NO KEY UPDATE". These don't block each
other, in contrast with already existing "SELECT FOR SHARE" and "SELECT
FOR UPDATE". UPDATE commands that do not modify the values stored in
the columns that are part of the key of the tuple now grab a SELECT FOR
NO KEY UPDATE lock on the tuple, allowing them to proceed concurrently
with tuple locks of the FOR KEY SHARE variety.
Foreign key triggers now use FOR KEY SHARE instead of FOR SHARE; this
means the concurrency improvement applies to them, which is the whole
point of this patch.
The added tuple lock semantics require some rejiggering of the multixact
module, so that the locking level that each transaction is holding can
be stored alongside its Xid. Also, multixacts now need to persist
across server restarts and crashes, because they can now represent not
only tuple locks, but also tuple updates. This means we need more
careful tracking of lifetime of pg_multixact SLRU files; since they now
persist longer, we require more infrastructure to figure out when they
can be removed. pg_upgrade also needs to be careful to copy
pg_multixact files over from the old server to the new, or at least part
of multixact.c state, depending on the versions of the old and new
servers.
Tuple time qualification rules (HeapTupleSatisfies routines) need to be
careful not to consider tuples with the "is multi" infomask bit set as
being only locked; they might need to look up MultiXact values (i.e.
possibly do pg_multixact I/O) to find out the Xid that updated a tuple,
whereas they previously were assured to only use information readily
available from the tuple header. This is considered acceptable, because
the extra I/O would involve cases that would previously cause some
commands to block waiting for concurrent transactions to finish.
Another important change is the fact that locking tuples that have
previously been updated causes the future versions to be marked as
locked, too; this is essential for correctness of foreign key checks.
This causes additional WAL-logging, also (there was previously a single
WAL record for a locked tuple; now there are as many as updated copies
of the tuple there exist.)
With all this in place, contention related to tuples being checked by
foreign key rules should be much reduced.
As a bonus, the old behavior that a subtransaction grabbing a stronger
tuple lock than the parent (sub)transaction held on a given tuple and
later aborting caused the weaker lock to be lost, has been fixed.
Many new spec files were added for isolation tester framework, to ensure
overall behavior is sane. There's probably room for several more tests.
There were several reviewers of this patch; in particular, Noah Misch
and Andres Freund spent considerable time in it. Original idea for the
patch came from Simon Riggs, after a problem report by Joel Jacobson.
Most code is from me, with contributions from Marti Raudsepp, Alexander
Shulgin, Noah Misch and Andres Freund.
This patch was discussed in several pgsql-hackers threads; the most
important start at the following message-ids:
AANLkTimo9XVcEzfiBR-ut3KVNDkjm2Vxh+t8kAmWjPuv@mail.gmail.com1290721684-sup-3951@alvh.no-ip.org1294953201-sup-2099@alvh.no-ip.org1320343602-sup-2290@alvh.no-ip.org1339690386-sup-8927@alvh.no-ip.org4FE5FF020200002500048A3D@gw.wicourts.gov4FEAB90A0200002500048B7D@gw.wicourts.gov
When a standby server follows the master using WAL archive, and it chooses
a new timeline (recovery_target_timeline='latest'), it only fetches the
timeline history file for the chosen target timeline, not any other history
files that might be missing from pg_xlog. For example, if the current
timeline is 2, and we choose 4 as the new recovery target timeline, the
history file for timeline 3 is not fetched, even if it's part of this
server's history. That's enough for the standby itself - the history file
for timeline 4 includes timeline 3 as well - but if a cascading standby
server wants to recover to timeline 3, it needs the history file. To fix,
when a new recovery target timeline is chosen, try to copy any missing
history files from the archive to pg_xlog between the old and new target
timeline.
A second similar issue was with the WAL files. When a standby recovers from
archive, and it reaches a segment that contains a switch to a new timeline,
recovery fetches only the WAL file labelled with the new timeline's ID. The
file from the new timeline contains a copy of the WAL from the old timeline
up to the point where the switch happened, and recovery recovers it from the
new file. But in streaming replication, walsender only tries to read it
from the old timeline's file. To fix, change walsender to read it from the
new file, so that it behaves the same as recovery in that sense, and doesn't
try to open the possibly nonexistent file with the old timeline's ID.
The xlogreader refactoring broke the logic to decide which timeline to start
streaming from. XLogPageRead() uses the timeline history to check which
timeline the requested WAL position falls into. However, after the
refactoring, XLogPageRead() is always first called with the first page in
the segment, to verify the segment header, and only then with the actual WAL
position we're interested in. That first read of the segment's header made
XLogPageRead() to always start streaming from the old timeline containing
the segment header, not the timeline containing the actual record, if there
was a timeline switch within the segment.
I thought I fixed this yesterday, but that fix was too narrow and only fixed
this for the corner-case that the timeline switch happened in the first page
of the segment. To fix this more robustly, pass explicitly the position of
the record we're actually interested in to XLogPageRead, and use that to
decide which timeline to read from, rather than deduce it from the page and
offset.
Per report from Fujii Masao.
get a large enough part of the page to include the beginning of the next
record we're interested in. The XLogPageRead callback uses the requested
length to decide which timeline to stream WAL from, and if the first call
is short, and the page contains a timeline switch, we'll repeatedly try
to stream that page from the old timeline, and never get across the
timeline switch.
This mirrors the changes done earlier to the server in standby mode. When
receivelog reaches the end of a timeline, as reported by the server, it
fetches the timeline history file of the next timeline, and restarts
streaming from the new timeline by issuing a new START_STREAMING command.
When pg_receivexlog crosses a timeline, it leaves the .partial suffix on the
last segment on the old timeline. This helps you to tell apart a partial
segment left in the directory because of a timeline switch, and a completed
segment. If you just follow a single server, it won't make a difference, but
it can be significant in more complicated scenarios where new WAL is still
generated on the old timeline.
This includes two small changes to the streaming replication protocol:
First, when you reach the end of timeline while streaming, the server now
sends the TLI of the next timeline in the server's history to the client.
pg_receivexlog uses that as the next timeline, so that it doesn't need to
parse the timeline history file like a standby server does. Second, when
BASE_BACKUP command sends the begin and end WAL positions, it now also sends
the timeline IDs corresponding the positions.
XLogReadRecord should reset its state on every error, to make sure it
re-reads the page on next call. It was inconsistent in that some errors did
that, but some did not.
In ReadRecord(), don't give up on an error if we're in standby mode. The
loop was set up to retry, but the checks within the loop broke out of the
loop on any error.
Andres Freund, with some tweaking by me.
The patch that turned XLogRecPtr into a uint64 inadvertently changed the
on-disk format of GiST indexes, because the NSN field in the GiST page
opaque is an XLogRecPtr. That breaks pg_upgrade. Revert the format of that
field back to the two-field struct that XLogRecPtr was before. This is the
same we did to LSNs in the page header to avoid changing on-disk format.
Bump catversion, as this invalidates any existing GiST indexes built on
9.3devel.
This new facility can not only be used by xlog.c to carry out crash
recovery, but also by external programs. By supplying a function to
read XLog pages from somewhere, all the WAL reading can be used for
completely different purposes.
For the standard backend use, the behavior should be pretty much the
same as previously. As for non-backend programs, an hypothetical
pg_xlogdump program is now closer to reality, but some more backend
support is still necessary.
This patch was originally submitted by Andres Freund in a different
form, but Heikki Linnakangas opted for and authored another design of
the concept. Andres has advanced the patch since Heikki's initial
version. Review and some (mostly cosmetics) changes by me.
Historically we've used a couple of very ad-hoc fudge factors to try to
get the right results when indexes of different sizes would satisfy a
query with the same number of index leaf tuples being visited. In
commit 21a39de580 I tweaked one of these
fudge factors, with results that proved disastrous for larger indexes.
Commit bf01e34b55 fudged it some more,
but still with not a lot of principle behind it.
What seems like a better way to address these issues is to explicitly model
index-descent costs, since that's what's really at stake when considering
diferent indexes with similar leaf-page-level costs. We tried that once
long ago, and found that charging random_page_cost per page descended
through was way too much, because upper btree levels tend to stay in cache
in real-world workloads. However, there's still CPU costs to think about,
and the previous fudge factors can be seen as a crude attempt to account
for those costs. So this patch replaces those fudge factors with explicit
charges for the number of tuple comparisons needed to descend the index
tree, plus a small charge per page touched in the descent. The cost
multipliers are chosen so that the resulting charges are in the vicinity of
the historical (pre-9.2) fudge factors for indexes of up to about a million
tuples, while not ballooning unreasonably beyond that, as the old fudge
factor did (even more so in 9.2).
To make this work accurately for btree indexes, add some code that allows
extraction of the known root-page height from a btree. There's no
equivalent number readily available for other index types, but we can use
the log of the number of index pages as an approximate substitute.
This seems like too much of a behavioral change to risk back-patching,
but it should improve matters going forward. In 9.2 I'll just revert
the fudge-factor change.
If you take a base backup from a standby server with "pg_basebackup -X
fetch", and the timeline switches while the backup is being taken, the
backup used to fail with an error "requested WAL segment %s has already
been removed". This is because the server-side code that sends over the
required WAL files would not construct the WAL filename with the correct
timeline after a switch.
Fix that by using readdir() to scan pg_xlog for all the WAL segments in the
range, regardless of timeline.
Also, include all timeline history files in the backup, if taken with
"-X fetch". That fixes another related bug: If a timeline switch happened
just before the backup was initiated in a standby, the WAL segment
containing the initial checkpoint record contains WAL from the older
timeline too. Recovery will not accept that without a timeline history file
that lists the older timeline.
Backpatch to 9.2. Versions prior to that were not affected as you could not
take a base backup from a standby before 9.2.
Streaming replication can fetch any missing timeline history files from the
master, but recovery would read the timeline history file for the target
timeline before reading the checkpoint record, and before walreceiver has
had a chance to fetch it from the master. Delay reading it, and the sanity
checks involving timeline history, until after reading the checkpoint
record.
There is at least one scenario where this makes a difference: if you take
a base backup from a standby server right after a timeline switch, the
WAL segment containing the initial checkpoint record will begin with an
older timeline ID. Without the timeline history file, recovering that file
will fail as the older timeline ID is not recognized to be an ancestor of
the target timeline. If you try to recover from such a backup, using only
streaming replication to fetch the WAL, this patch is required for that to
work.
After receiving some WAL over streaming replication, try to open the file
from the timeline we're currently recieving, not recoveryTargetTLI. They
are usually the same, which is why wasn't noticed before, but you'd get
an error if there have been more than one timeline switch between the
current point in WAL and the recovery target.
The cascading standby patch in 9.2 changed the way WAL files are treated
when restored from the archive. Before, they were restored under a temporary
filename, and not kept in pg_xlog, but after the patch, they were copied
under pg_xlog. This is necessary for a cascading standby to find them, but
it also means that if the archive goes offline and a standby is restarted,
it can recover back to where it was using the files in pg_xlog. It also
means that if you take an offline backup from a standby server, it includes
all the required WAL files in pg_xlog.
However, the same change was not made to timeline history files, so if the
WAL segment containing the checkpoint record contains a timeline switch, you
will still get an error if you try to restart recovery without the archive,
or recover from an offline backup taken from the standby.
With this patch, timeline history files restored from archive are copied
into pg_xlog like WAL files are, so that pg_xlog contains all the files
required to recover. This is a corner-case pre-existing issue in 9.2, but
even more important in master where it's possible for a standby to follow a
timeline switch through streaming replication. To make that possible, the
timeline history files must be present in pg_xlog.
This gets rid of XLByteLT, XLByteLE, XLByteEQ and XLByteAdvance.
These were useful for brevity when XLogRecPtrs were split in
xlogid/xrecoff; but now that they are simple uint64's, they are just
clutter. The only downside to making this change would be ease of
backporting patches, but that has been negated by other substantive
changes to the involved code anyway. The clarity of simpler expressions
makes the change worthwhile.
Most of the changes are mechanical, but in a couple of places, the patch
author chose to invert the operator sense, making the code flow more
logical (and more in line with preceding comments).
Author: Andres Freund
Eyeballed by Dimitri Fontaine and Alvaro Herrera
This was broken before, we would recycle old WAL segments on wrong timeline
after the recovery target timeline had changed, but my recent commit to
not initialize ThisTimeLineID at all in a standby's checkpointer process
broke this completely.
The problem is that when installing a recycled WAL segment as a future one,
ThisTimeLineID is used to construct the filename. To fix, always update
ThisTimeLineID to the current timeline being recovered, before recycling
WAL segments at a restartpoint.
This still leaves a small window where we might install WAL segments under
wrong timeline ID, if the timeline is changed just as we're about to start
recycling. Also, even if we're replaying timeline X at the momnent, there's
no guarantee that we'll need as many WAL segments on that timeline as we
recycle. We might be just about to reach the point where we switch to next
timeline, so might only need one more WAL segment on the current timeline.
We'll live with the waste in that situation.
Bug pointed out by Fujii Masao. 9.1 and 9.2 had the same issue, when
recovery target timeline was changed, but I committed a slightly different
version of this patch on those branches.
Most of the time, the last replayed record comes from the recovery target
timeline, but there is a corner case where it makes a difference. When
the startup process scans for a new timeline, and decides to change recovery
target timeline, there is a window where the recovery target TLI has already
been bumped, but there are no WAL segments from the new timeline in pg_xlog
yet. For example, if we have just replayed up to point 0/30002D8, on
timeline 1, there is a WAL file called 000000010000000000000003 in pg_xlog
that contains the WAL up to that point. When recovery switches recovery
target timeline to 2, a walsender can immediately try to read WAL from
0/30002D8, from timeline 2, so it will try to open WAL file
000000020000000000000003. However, that doesn't exist yet - the startup
process hasn't copied that file from the archive yet nor has the walreceiver
streamed it yet, so walsender fails with error "requested WAL segment
000000020000000000000003 has already been removed". That's harmless, in that
the standby will try to reconnect later and by that time the segment is
already created, but error messages that should be ignored are not good.
To fix that, have walsender track the TLI of the last replayed record,
instead of the recovery target timeline. That way walsender will not try to
read anything from timeline 2, until the WAL segment has been created and at
least one record has been replayed from it. The recovery target timeline is
now xlog.c's internal affair, it doesn't need to be exposed in shared memory
anymore.
This fixes the error reported by Thom Brown. depesz the same error message,
but I'm not sure if this fixes his scenario.
We used to set it to the current recovery target timeline, but the recovery
target timeline can change during recovery, leaving ThisTimeLineID at an
old value. That seems worse than always leaving it at zero to begin with.
AFAICS there was no good reason to set it in the first place. ThisTimeLineID
is not needed in checkpointer or bgwriter process, until it's time to write
the end-of-recovery checkpoint, and at that point ThisTimeLineID is updated
anyway.
Heikki Linnakangas
Robert Haas
Tom Lane
Alvaro Herrera
Bruce Momjian
Magnus Hagander
Peter Eisentraut
Fujii Masao
Simon Riggs
Noah Misch
Jeff Davis
Greg Stark
Stephen Frost
Kevin Grittner