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These were missed in commit a213f1ee6c5a1bbe1f074ca201975e76ad2ed50c,
which removed that function.
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SP-GiST is comparable to GiST in flexibility, but supports non-balanced
partitioned search structures rather than balanced trees. As described at
PGCon 2011, this new indexing structure can beat GiST in both index build
time and query speed for search problems that it is well matched to.
There are a number of areas that could still use improvement, but at this
point the code seems committable.
Teodor Sigaev and Oleg Bartunov, with considerable revisions by Tom Lane
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Removing this bit from xl_info allows us to restore the old limit of four
(not three) separate pages touched by a WAL record, which is needed for the
upcoming SP-GiST feature, and will likely be useful elsewhere in future.
When we implemented XLR_BKP_REMOVABLE in 2007, we had to do it like that
because no special WAL-visible action was taken when starting a backup.
However, now we force a segment switch when starting a backup, so a
compressing WAL archiver (such as pglesslog) that uses the state shown in
the current page header will not be fooled as to removability of backup
blocks. The only downside is that the archiver will not return to
compressing mode for up to one WAL page after the backup is over, which is
a small price to pay for getting back the extra xl_info bit. In any case
the archiver could look for XLOG_BACKUP_END records if it thought it was
worth the trouble to do so.
Bump XLOG_PAGE_MAGIC since this is effectively a change in WAL format.
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we don't reach consistency before replaying all of the WAL. Rename the
variable to reachedConsistency, to make its intention clearer.
In master, that was an active bug because of the recent patch to
immediately PANIC if a reference to a missing page is found in WAL after
reaching consistency, as Tom Lane's test case demonstrated. In 9.1 and 9.0,
the only consequence was a misleading "consistent recovery state reached at
%X/%X" message in the log at the beginning of crash recovery (the database
is not consistent at that point yet). In 8.4, the log message was not
printed in crash recovery, even though there was a similar
reachedMinRecoveryPoint local variable that was also set early. So,
backpatch to 9.1 and 9.0.
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This patch creates an API whereby a btree index opclass can optionally
provide non-SQL-callable support functions for sorting. In the initial
patch, we only use this to provide a directly-callable comparator function,
which can be invoked with a bit less overhead than the traditional
SQL-callable comparator. While that should be of value in itself, the real
reason for doing this is to provide a datatype-extensible framework for
more aggressive optimizations, as in Peter Geoghegan's recent work.
Robert Haas and Tom Lane
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invalid-page hash table, PANIC immediately. Immediate PANIC is much better
than waiting for end-of-recovery, which is what we did before, because the
end-of-recovery might not come until months later if this is a standby
server.
Also refrain from creating a restartpoint if there are invalid-page entries
in the hash table. Restarting recovery from such a restartpoint would not
see the invalid references, and wouldn't be able to cross-check them when
consistency is reached. That wouldn't matter when things are going smoothly,
but the more sanity checks you have the better.
Fujii Masao
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In the previous coding, callers were faced with an awkward choice:
look up the name, do permissions checks, and then lock the table; or
look up the name, lock the table, and then do permissions checks.
The first choice was wrong because the results of the name lookup
and permissions checks might be out-of-date by the time the table
lock was acquired, while the second allowed a user with no privileges
to interfere with access to a table by users who do have privileges
(e.g. if a malicious backend queues up for an AccessExclusiveLock on
a table on which AccessShareLock is already held, further attempts
to access the table will be blocked until the AccessExclusiveLock
is obtained and the malicious backend's transaction rolls back).
To fix, allow callers of RangeVarGetRelid() to pass a callback which
gets executed after performing the name lookup but before acquiring
the relation lock. If the name lookup is retried (because
invalidation messages are received), the callback will be re-executed
as well, so we get the best of both worlds. RangeVarGetRelid() is
renamed to RangeVarGetRelidExtended(); callers not wishing to supply
a callback can continue to invoke it as RangeVarGetRelid(), which is
now a macro. Since the only one caller that uses nowait = true now
passes a callback anyway, the RangeVarGetRelid() macro defaults nowait
as well. The callback can also be used for supplemental locking - for
example, REINDEX INDEX needs to acquire the table lock before the index
lock to reduce deadlock possibilities.
There's a lot more work to be done here to fix all the cases where this
can be a problem, but this commit provides the general infrastructure
and fixes the following specific cases: REINDEX INDEX, REINDEX TABLE,
LOCK TABLE, and and DROP TABLE/INDEX/SEQUENCE/VIEW/FOREIGN TABLE.
Per discussion with Noah Misch and Alvaro Herrera.
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Use of a randomly chosen large value was never exactly graceful, and
now that there are penalty functions that are intentionally using infinity,
it doesn't seem like a good idea for null-vs-not-null to be using something
less.
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Jeff Janes
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A simple thinko in ginRedoUpdateMetapage, namely failing to increment a
loop counter, led to inserting records into the last pending-list page in
the wrong order (the opposite of that intended). So far as I can tell,
this would not upset the code that eventually flushes pending items into
the main part of the GIN index. But it did break the code that searched
the pending list for matches, resulting in transient failure to find
matching entries during index lookups, as illustrated in bug #6307 from
Maksym Boguk.
Back-patch to 8.4 where the incorrect code was introduced.
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This speeds up snapshot-taking and reduces ProcArrayLock contention.
Also, the PGPROC (and PGXACT) structures used by two-phase commit are
now allocated as part of the main array, rather than in a separate
array, and we keep ProcArray sorted in pointer order. These changes
are intended to minimize the number of cache lines that must be pulled
in to take a snapshot, and testing shows a substantial increase in
performance on both read and write workloads at high concurrencies.
Pavan Deolasee, Heikki Linnakangas, Robert Haas
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even when there is no work to do. Further analysis required.
Revert of patch c1458cc495ff800cd176a1c2e56d8b62680d9b71
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When wal_level = 'hot_standby' we touched the last page of the
relation during a VACUUM, even if nothing else had happened.
That would alter the LSN of the last block and set the mtime
of the relation file unnecessarily. Noted by Thom Brown.
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Previously we waited for wal_writer_delay before flushing WAL. Now
we also wake WALWriter as soon as a WAL buffer page has filled.
Significant effect observed on performance of asynchronous commits
by Robert Haas, attributed to the ability to set hint bits on tuples
earlier and so reducing contention caused by clog lookups.
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I got alignment wrong in the redo routine. Spotted by redoing the log
genereated by copy regression test.
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Forgot to call RestoreBkpBlocks() in the redo-function, as pointed out by
Simon Riggs. In redo of a regular heap insert, it's taken care of in
heap_redo(), but this new record type uses the heap2 RM, and heap2_redo()
does not take care of that for you.
Also, failed to reset the vmbuffer and all_visibile_cleared local variables
after switching to a new buffer.
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This greatly reduces the WAL volume, especially when the table is narrow.
The overhead of locking the heap page is also reduced. Reduced WAL traffic
also makes it scale a lot better, if you run multiple COPY processes at
the same time.
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In a regular VACUUM, it's OK to skip pages for which a cleanup lock
isn't immediately available; the next VACUUM will deal with them. If
we're scanning the entire relation to advance relfrozenxid, we might
need to wait, but only if there are tuples on the page that actually
require freezing. These changes should greatly reduce the incidence
of of vacuum processes getting "stuck".
Simon Riggs and Robert Haas
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This assumption can be wrong when the toaster is passed a raw on-disk
tuple, because the tuple might pre-date an ALTER TABLE ADD COLUMN operation
that added columns without rewriting the table. In such a case the tuple's
natts value is smaller than what we expect from the tuple descriptor, and
so its t_hoff value could be smaller too. In fact, the tuple might not
have a null bitmap at all, and yet our current opinion of it is that it
contains some trailing nulls.
In such a situation, toast_insert_or_update did the wrong thing, because
to save a few lines of code it would use the old t_hoff value as the offset
where heap_fill_tuple should start filling data. This did not leave enough
room for the new nulls bitmap, with the result that the first few bytes of
data could be overwritten with null flag bits, as in a recent report from
Hubert Depesz Lubaczewski.
The particular case reported requires ALTER TABLE ADD COLUMN followed by
CREATE TABLE AS SELECT * FROM ... or INSERT ... SELECT * FROM ..., and
further requires that there be some out-of-line toasted fields in one of
the tuples to be copied; else we'll not reach the troublesome code.
The problem can only manifest in this form in 8.4 and later, because
before commit a77eaa6a95009a3441e0d475d1980259d45da072, CREATE TABLE AS or
INSERT/SELECT wouldn't result in raw disk tuples getting passed directly
to heap_insert --- there would always have been at least a junkfilter in
between, and that would reconstitute the tuple header with an up-to-date
t_natts and hence t_hoff. But I'm backpatching the tuptoaster change all
the way anyway, because I'm not convinced there are no older code paths
that present a similar risk.
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If we have an inequality key that constrains the other end of the index,
it doesn't directly help us in doing the initial positioning ... but it
does imply a NOT NULL constraint on the index column. If the index stores
nulls at this end, we can use the implied NOT NULL condition for initial
positioning, just as if it had been stated explicitly. This avoids wasting
time when there are a lot of nulls in the column. This is the reverse of
the examples given in bugs #6278 and #6283, which were about failing to
stop early when we encounter nulls at the end of the indexscan.
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As pointed out by Naoya Anzai, my previous try at this was a few bricks
shy of a load, because I had forgotten that the initial-positioning logic
might not try to skip over nulls at the end of the index the scan will
start from. We ought to fix that, because it represents an unnecessary
inefficiency, but first let's get the scan-stop logic back to a safe
state. With this patch, we preserve the performance benefit requested
in bug #6278 for the case of scanning forward into NULLs (in a NULLS
LAST index), but the reverse case of scanning backward across NULLs
when there's no suitable initial-positioning qual is still inefficient.
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Previously, we skipped a checkpoint if no WAL had been written since
last checkpoint, though this does not appear in user documentation.
As of now, we skip a checkpoint until we have written at least one
enough WAL to switch the next WAL file. This greatly reduces the
level of activity and number of WAL messages generated by a very
low activity server. This is safe because the purpose of a checkpoint
is to act as a starting place for a recovery, in case of crash.
This patch maintains minimal WAL volume for replay in case of crash,
thus maintaining very low crash recovery time.
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Startup process now has its own dedicated file, just like all other
special/background processes. Reduces role and size of xlog.c
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There was a timing window between when oldestActiveXid was derived
and when it should have been derived that only shows itself under
heavy load. Move code around to ensure correct timing of derivation.
No change to StartupSUBTRANS() code, which is where this failed.
Bug report by Chris Redekop
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Patch by me, bug report by Chris Redekop, analysis by Florian Pflug
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If a tuple in a syscache contains an out-of-line toasted field, and we
try to fetch that field shortly after some other transaction has committed
an update or deletion of the tuple, there is a race condition: vacuum
could come along and remove the toast tuples before we can fetch them.
This leads to transient failures like "missing chunk number 0 for toast
value NNNNN in pg_toast_2619", as seen in recent reports from Andrew
Hammond and Tim Uckun.
The design idea of syscache is that access to stale syscache entries
should be prevented by relation-level locks, but that fails for at least
two cases where toasted fields are possible: ANALYZE updates pg_statistic
rows without locking out sessions that might want to plan queries on the
same table, and CREATE OR REPLACE FUNCTION updates pg_proc rows without
any meaningful lock at all.
The least risky fix seems to be an idea that Heikki suggested when we
were dealing with a related problem back in August: forcibly detoast any
out-of-line fields before putting a tuple into syscache in the first place.
This avoids the problem because at the time we fetch the parent tuple from
the catalog, we should be holding an MVCC snapshot that will prevent
removal of the toast tuples, even if the parent tuple is outdated
immediately after we fetch it. (Note: I'm not convinced that this
statement holds true at every instant where we could be fetching a syscache
entry at all, but it does appear to hold true at the times where we could
fetch an entry that could have a toasted field. We will need to be a bit
wary of adding toast tables to low-level catalogs that don't have them
already.) An additional benefit is that subsequent uses of the syscache
entry should be faster, since they won't have to detoast the field.
Back-patch to all supported versions. The problem is significantly harder
to reproduce in pre-9.0 releases, because of their willingness to flush
every entry in a syscache whenever the underlying catalog is vacuumed
(cf CatalogCacheFlushRelation); but there is still a window for trouble.
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The existing scan-direction-sensitive tests were overly complex, and
failed to stop the scan in cases where it's perfectly legitimate to do so.
Per bug #6278 from Maksym Boguk.
Back-patch to 8.3, which is as far back as the patch applies easily.
Doesn't seem worth sweating over a relatively minor performance issue in
8.2 at this late date. (But note that this was a performance regression
from 8.1 and before, so 8.2 is being left as an outlier.)
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Recent work on index-only scans left this somewhat out of date.
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A transaction can export a snapshot with pg_export_snapshot(), and then
others can import it with SET TRANSACTION SNAPSHOT. The data does not
leave the server so there are not security issues. A snapshot can only
be imported while the exporting transaction is still running, and there
are some other restrictions.
I'm not totally convinced that we've covered all the bases for SSI (true
serializable) mode, but it works fine for lesser isolation modes.
Joachim Wieland, reviewed by Marko Tiikkaja, and rather heavily modified
by Tom Lane
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This is merely an exercise in satisfying pedants, not a bug fix, because
in every case we were checking for failure later with ferror(), or else
there was nothing useful to be done about a failure anyway. Document
the latter cases.
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In general the data returned by an index-only scan should have the
datatypes originally computed by FormIndexDatum. If the index opclasses
use "storage" datatypes different from their input datatypes, the scan
tuple will not have the same rowtype attributed to the index; but we had
a hard-wired assumption that that was true in nodeIndexonlyscan.c. We'd
already hacked around the issue for the one case where the types are
different in btree indexes (btree name_ops), but this would definitely
come back to bite us if we ever implement index-only scans in GiST.
To fix, require the index AM to explicitly provide the tupdesc for the
tuple it is returning. btree can just pass back the index's tupdesc, but
GiST will have to work harder when and if it supports index-only scans.
I had previously proposed fixing this by allowing the index AM to fill the
scan tuple slot directly; but on reflection that seemed like a module
layering violation, since TupleTableSlots are creatures of the executor.
At least in the btree case, it would also be less efficient, since the
tuple deconstruction work would occur even for rows later found to be
invisible to the scan's snapshot.
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This allows "indexedcol op ANY(ARRAY[...])" conditions to be used in plain
indexscans, and particularly in index-only scans.
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Add a column pg_class.relallvisible to remember the number of pages that
were all-visible according to the visibility map as of the last VACUUM
(or ANALYZE, or some other operations that update pg_class.relpages).
Use relallvisible/relpages, instead of an arbitrary constant, to estimate
how many heap page fetches can be avoided during an index-only scan.
This is pretty primitive and will no doubt see refinements once we've
acquired more field experience with the index-only scan mechanism, but
it's way better than using a constant.
Note: I had to adjust an underspecified query in the window.sql regression
test, because it was changing answers when the plan changed to use an
index-only scan. Some of the adjacent tests perhaps should be adjusted
as well, but I didn't do that here.
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struct, to help pgindent.
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It's been bothering me for several days that pretending that the cstring
data stored in a btree name_ops column is really a "name" Datum could lead
to reading past the end of memory. However, given the current memory
layout used for index-only scans in the btree code, a crash is in fact not
possible. Document that so we don't break it. I have not thought of any
other solutions that aren't fairly ugly too, and most of them lose the
functionality of index-only scans on name columns altogether, so this seems
like the way to go.
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The original idea of this patch was to make box picksplit run faster, by
eliminating unnecessary palloc() overhead, but that was obsoleted by the new
double-sorting split algorithm that doesn't call these functions so heavily
anymore. Nevertheless, the code looks better this way.
Original patch by me, reviewed and tidied up after the double-sorting patch
by Kevin Grittner.
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We copy all the matched tuples off the page during _bt_readpage, instead of
expensively re-locking the page during each subsequent tuple fetch. This
costs a bit more local storage, but not more than 2*BLCKSZ worth, and the
reduction in LWLock traffic is certainly worth that. What's more, this
lets us get rid of the API wart in the original patch that said an index AM
could randomly decline to supply an index tuple despite having asserted
pg_am.amcanreturn. That will be important for future improvements in the
index-only-scan feature, since the executor will now be able to rely on
having the index data available.
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When a btree index contains all columns required by the query, and the
visibility map shows that all tuples on a target heap page are
visible-to-all, we don't need to fetch that heap page. This patch depends
on the previous patches that made the visibility map reliable.
There's a fair amount left to do here, notably trying to figure out a less
chintzy way of estimating the cost of an index-only scan, but the core
functionality seems ready to commit.
Robert Haas and Ibrar Ahmed, with some previous work by Heikki Linnakangas.
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new double-sorting algorithm. The new algorithm produces better quality
trees, making searches faster.
Alexander Korotkov
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Previously, the code assumed that the only possible action to take was
to delete files behind a certain cutoff point. The async notify code
was already a crock: it used a different "pagePrecedes" function for
truncation than for regular operation. By allowing it to pass a
callback to SlruScanDirectory it can do cleanly exactly what it needs to
do.
The clog.c code also had its own use for SlruScanDirectory, which is
made a bit simpler with this.
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This patch has two distinct purposes: to report multiple problems in
postgresql.conf rather than always bailing out after the first one,
and to change the policy for whether changes are applied when there are
unrelated errors in postgresql.conf.
Formerly the policy was to apply no changes if any errors could be
detected, but that had a significant consistency problem, because in some
cases specific values might be seen as valid by some processes but invalid
by others. This meant that the latter processes would fail to adopt
changes in other parameters even though the former processes had done so.
The new policy is that during SIGHUP, the file is rejected as a whole
if there are any errors in the "name = value" syntax, or if any lines
attempt to set nonexistent built-in parameters, or if any lines attempt
to set custom parameters whose prefix is not listed in (the new value of)
custom_variable_classes. These tests should always give the same results
in all processes, and provide what seems a reasonably robust defense
against loading values from badly corrupted config files. If these tests
pass, all processes will apply all settings that they individually see as
good, ignoring (but logging) any they don't.
In addition, the postmaster does not abandon reading a configuration file
after the first syntax error, but continues to read the file and report
syntax errors (up to a maximum of 100 syntax errors per file).
The postmaster will still refuse to start up if the configuration file
contains any errors at startup time, but these changes allow multiple
errors to be detected and reported before quitting.
Alexey Klyukin, reviewed by Andy Colson and av (Alexander ?)
with some additional hacking by Tom Lane
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pg_trgm was already doing this unofficially, but the implementation hadn't
been thought through very well and leaked memory. Restructure the core
GiST code so that it actually works, and document it. Ordinarily this
would have required an extra memory context creation/destruction for each
GiST index search, but I was able to avoid that in the normal case of a
non-rescanned search by finessing the handling of the RBTree. It used to
have its own context always, but now shares a context with the
scan-lifespan data structures, unless there is more than one rescan call.
This should make the added overhead unnoticeable in typical cases.
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In hio.c, document how we avoid deadlock with respect to visibility map
buffer locks. In visibilitymap.c, update the LOCKING section of the
file header comment.
Both oversights noted by Heikki Linnakangas.
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Bug found by Alvaro Herrera, fix suggested by Heikki Linnakangas
and reviewed by Tom Lane.
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In REPEATABLE READ (nee SERIALIZABLE) mode, an attempt to do
GetTransactionSnapshot() between AbortTransaction and CleanupTransaction
failed, because GetTransactionSnapshot would recompute the transaction
snapshot (which is already wrong, given the isolation mode) and then
re-register it in the TopTransactionResourceOwner, leading to an Assert
because the TopTransactionResourceOwner should be empty of resources after
AbortTransaction. This is the root cause of bug #6218 from Yamamoto
Takashi. While changing plancache.c to avoid requesting a snapshot when
handling a ROLLBACK masks the problem, I think this is really a snapmgr.c
bug: it's lower-level than the resource manager mechanism and should not be
shutting itself down before we unwind resource manager resources. However,
just postponing the release of the transaction snapshot until cleanup time
didn't work because of the circular dependency with
TopTransactionResourceOwner. Fix by managing the internal reference to
that snapshot manually instead of depending on TopTransactionResourceOwner.
This saves a few cycles as well as making the module layering more
straightforward. predicate.c's dependencies on TopTransactionResourceOwner
go away too.
I think this is a longstanding bug, but there's no evidence that it's more
than a latent bug, so it doesn't seem worth any risk of back-patching.
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