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.
Before this patch, streaming replication would refuse to start replicating
if the timeline in the primary doesn't exactly match the standby. The
situation where it doesn't match is when you have a master, and two
standbys, and you promote one of the standbys to become new master.
Promoting bumps up the timeline ID, and after that bump, the other standby
would refuse to continue.
There's significantly more timeline related logic in streaming replication
now. First of all, when a standby connects to primary, it will ask the
primary for any timeline history files that are missing from the standby.
The missing files are sent using a new replication command TIMELINE_HISTORY,
and stored in standby's pg_xlog directory. Using the timeline history files,
the standby can follow the latest timeline present in the primary
(recovery_target_timeline='latest'), just as it can follow new timelines
appearing in an archive directory.
START_REPLICATION now takes a TIMELINE parameter, to specify exactly which
timeline to stream WAL from. This allows the standby to request the primary
to send over WAL that precedes the promotion. The replication protocol is
changed slightly (in a backwards-compatible way although there's little hope
of streaming replication working across major versions anyway), to allow
replication to stop when the end of timeline reached, putting the walsender
back into accepting a replication command.
Many thanks to Amit Kapila for testing and reviewing various versions of
this patch.
This allows us to do some more rigorous sanity checking for various
incorrect point-in-time recovery scenarios, and provides more information
for debugging purposes. It will also come handy in the upcoming patch to
allow timeline switches to be replicated by streaming replication.
Allow support only for freezing tuples by explicit
command. Previous coding mistakenly extended
slightly beyond what was agreed as correct on -hackers.
So essentially a partial revoke of earlier work,
leaving just the COPY FREEZE command.
When a relfilenode is created in this subtransaction or
a committed child transaction and it cannot otherwise
be seen by our own process, mark tuples committed ahead
of transaction commit for all COPY commands in same
transaction. If FREEZE specified on COPY
and pre-conditions met then rows will also be frozen.
Both options designed to avoid revisiting rows after commit,
increasing performance of subsequent commands after
data load and upgrade. pg_restore changes later.
Simon Riggs, review comments from Heikki Linnakangas, Noah Misch and design
input from Tom Lane, Robert Haas and Kevin Grittner
When I moved ExecuteRecoveryCommand() from xlog.c to xlogarchive.c, I didn't
realize that it's called from the checkpoint process, not the startup
process. I tried to use InRedo variable to decide whether or not to attempt
cleaning up the archive (must not do so before we have read the initial
checkpoint record), but that variable is only valid within the startup
process.
Instead, let ExecuteRecoveryCommand() always clean up the archive, and add
an explicit argument to RestoreArchivedFile() to say whether that's allowed
or not. The caller knows better.
Reported by Erik Rijkers, diagnosis by Fujii Masao. Only 9.3devel is
affected.
Most of the replay functions for WAL record types that modify more than
one page failed to ensure that those pages were locked correctly to ensure
that concurrent queries could not see inconsistent page states. This is
a hangover from coding decisions made long before Hot Standby was added,
when it was hardly necessary to acquire buffer locks during WAL replay
at all, let alone hold them for carefully-chosen periods.
The key problem was that RestoreBkpBlocks was written to hold lock on each
page restored from a full-page image for only as long as it took to update
that page. This was guaranteed to break any WAL replay function in which
there was any update-ordering constraint between pages, because even if the
nominal order of the pages is the right one, any mixture of full-page and
non-full-page updates in the same record would result in out-of-order
updates. Moreover, it wouldn't work for situations where there's a
requirement to maintain lock on one page while updating another. Failure
to honor an update ordering constraint in this way is thought to be the
cause of bug #7648 from Daniel Farina: what seems to have happened there
is that a btree page being split was rewritten from a full-page image
before the new right sibling page was written, and because lock on the
original page was not maintained it was possible for hot standby queries to
try to traverse the page's right-link to the not-yet-existing sibling page.
To fix, get rid of RestoreBkpBlocks as such, and instead create a new
function RestoreBackupBlock that restores just one full-page image at a
time. This function can be invoked by WAL replay functions at the points
where they would otherwise perform non-full-page updates; in this way, the
physical order of page updates remains the same no matter which pages are
replaced by full-page images. We can then further adjust the logic in
individual replay functions if it is necessary to hold buffer locks
for overlapping periods. A side benefit is that we can simplify the
handling of concurrency conflict resolution by moving that code into the
record-type-specfic functions; there's no more need to contort the code
layout to keep conflict resolution in front of the RestoreBkpBlocks call.
In connection with that, standardize on zero-based numbering rather than
one-based numbering for referencing the full-page images. In HEAD, I
removed the macros XLR_BKP_BLOCK_1 through XLR_BKP_BLOCK_4. They are
still there in the header files in previous branches, but are no longer
used by the code.
In addition, fix some other bugs identified in the course of making these
changes:
spgRedoAddNode could fail to update the parent downlink at all, if the
parent tuple is in the same page as either the old or new split tuple and
we're not doing a full-page image: it would get fooled by the LSN having
been advanced already. This would result in permanent index corruption,
not just transient failure of concurrent queries.
Also, ginHeapTupleFastInsert's "merge lists" case failed to mark the old
tail page as a candidate for a full-page image; in the worst case this
could result in torn-page corruption.
heap_xlog_freeze() was inconsistent about using a cleanup lock or plain
exclusive lock: it did the former in the normal path but the latter for a
full-page image. A plain exclusive lock seems sufficient, so change to
that.
Also, remove gistRedoPageDeleteRecord(), which has been dead code since
VACUUM FULL was rewritten.
Back-patch to 9.0, where hot standby was introduced. Note however that 9.0
had a significantly different WAL-logging scheme for GIST index updates,
and it doesn't appear possible to make that scheme safe for concurrent hot
standby queries, because it can leave inconsistent states in the index even
between WAL records. Given the lack of complaints from the field, we won't
work too hard on fixing that branch.
This prevents surprising behavior when a FOR EACH ROW trigger
BEFORE UPDATE or BEFORE DELETE directly or indirectly updates or
deletes the the old row. Prior to this patch the requested action
on the row could be silently ignored while all triggered actions
based on the occurence of the requested action could be committed.
One example of how this could happen is if the BEFORE DELETE
trigger for a "parent" row deleted "children" which had trigger
functions to update summary or status data on the parent.
This also prevents similar surprising problems if the query has a
volatile function which updates a target row while it is already
being updated.
There are related issues present in FOR UPDATE cursors and READ
COMMITTED queries which are not handled by this patch. These
issues need further evalution to determine what change, if any, is
needed.
Where the new error messages are generated, in most cases the best
fix will be to move code from the BEFORE trigger to an AFTER
trigger. Where this is not feasible, the trigger can avoid the
error by re-issuing the triggering statement and returning NULL.
Documentation changes will be submitted in a separate patch.
Kevin Grittner and Tom Lane with input from Florian Pflug and
Robert Haas, based on problems encountered during conversion of
Wisconsin Circuit Court trigger logic to plpgsql triggers.
This is just refactoring, to make the functions accessible outside xlog.c.
A followup patch will make use of that, to allow fetching timeline history
files over streaming replication.
This fixes another error in commit 9e8da0f75731aaa7605cf4656c21ea09e84d2eb1.
I neglected to make the mark/restore functionality save and restore the
current set of array key values, which led to strange behavior if an
IndexScan with ScalarArrayOpExpr quals was used as the inner side of a
mergejoin. Per bug #7570 from Melese Tesfaye.
The cascading replication code assumed that the current RecoveryTargetTLI
never changes, but that's not true with recovery_target_timeline='latest'.
The obvious upshot of that is that RecoveryTargetTLI in shared memory needs
to be protected by a lock. A less obvious consequence is that when a
cascading standby is connected, and the standby switches to a new target
timeline after scanning the archive, it will continue to stream WAL to the
cascading standby, but from a wrong file, ie. the file of the previous
timeline. For example, if the standby is currently streaming from the middle
of file 000000010000000000000005, and the timeline changes, the standby
will continue to stream from that file. However, the WAL on the new
timeline is in file 000000020000000000000005, so the standby sends garbage
from 000000010000000000000005 to the cascading standby, instead of the
correct WAL from file 000000020000000000000005.
This also fixes a related bug where a partial WAL segment is restored from
the archive and streamed to a cascading standby. The code assumed that when
a WAL segment is copied from the archive, it can immediately be fully
streamed to a cascading standby. However, if the segment is only partially
filled, ie. has the right size, but only N first bytes contain valid WAL,
that's not safe. That can happen if a partial WAL segment is manually copied
to the archive, or if a partial WAL segment is archived because a server is
started up on a new timeline within that segment. The cascading standby will
get confused if the WAL it received is not valid, and will get stuck until
it's restarted. This patch fixes that problem by not allowing WAL restored
from the archive to be streamed to a cascading standby until it's been
replayed, and thus validated.
This reduces unnecessary exposure of other headers through htup.h, which
is very widely included by many files.
I have chosen to move the function prototypes to the new file as well,
because that means htup.h no longer needs to include tupdesc.h. In
itself this doesn't have much effect in indirect inclusion of tupdesc.h
throughout the tree, because it's also required by execnodes.h; but it's
something to explore in the future, and it seemed best to do the htup.h
change now while I'm busy with it.
The heapam XLog functions are used by other modules, not all of which
are interested in the rest of the heapam API. With this, we let them
get just the XLog stuff in which they are interested and not pollute
them with unrelated includes.
Also, since heapam.h no longer requires xlog.h, many files that do
include heapam.h no longer get xlog.h automatically, including a few
headers. This is useful because heapam.h is getting pulled in by
execnodes.h, which is in turn included by a lot of files.
We use a hash table to track the parents of inner pages, but when inserting
to a leaf page, the caller of gistbufferinginserttuples() must pass a
correct block number of the leaf's parent page. Before gistProcessItup()
descends to a child page, it checks if the downlink needs to be adjusted to
accommodate the new tuple, and updates the downlink if necessary. However,
updating the downlink might require splitting the page, which might move the
downlink to a page to the right. gistProcessItup() doesn't realize that, so
when it descends to the leaf page, it might pass an out-of-date parent block
number as a result. Fix that by returning the block a tuple was inserted to
from gistbufferinginserttuples().
This fixes the bug reported by Zdeněk Jílovec.
The correct test for whether a redirection tuple is removable is whether
tuple's xid < RecentGlobalXmin, not OldestXmin; the previous coding
failed to protect index searches being done in concurrent transactions that
have no XID. This mirrors the recent fix in btree's page recycling logic
made in commit d3abbbebe52eb1e59e621c880ad57df9d40d13f2.
Also, WAL-log the newest XID of any removed redirection tuple on an index
page, and apply ResolveRecoveryConflictWithSnapshot during InHotStandby WAL
replay. This protects against concurrent Hot Standby transactions possibly
needing to see the redirection tuple(s).
Per my query of 2012-03-12 and subsequent discussion.
Commit 061e7efb1b4c5b8a5d02122b7780531b8d5bf23d changed the rules
for splitting xlog records across pages, but neglected to update this
test. It's possible that there's some better action here than just
removing the test completely, but this at least appears to get some
of the things that are currently broken (like initdb on MacOS X)
working again.
The latter was already the dominant use, and it's preferable because
in C the convention is that intXX means XX bits. Therefore, allowing
mixed use of int2, int4, int8, int16, int32 is obviously confusing.
Remove the typedefs for int2 and int4 for now. They don't seem to be
widely used outside of the PostgreSQL source tree, and the few uses
can probably be cleaned up by the time this ships.
This simplifies code that needs to do arithmetic on XLogRecPtrs.
To avoid changing on-disk format of data pages, the LSN on data pages is
still stored in the old format. That should keep pg_upgrade happy. However,
we have XLogRecPtrs embedded in the control file, and in the structs that
are sent over the replication protocol, so this changes breaks compatibility
of pg_basebackup and server. I didn't do anything about this in this patch,
per discussion on -hackers, the right thing to do would to be to change the
replication protocol to be architecture-independent, so that you could use
a newer version of pg_receivexlog, for example, against an older server
version.
This saves a few bytes of WAL space, but the real motivation is to make it
predictable how much WAL space a record requires, as it no longer depends
on whether we need to waste the last few bytes at end of WAL page because
the header doesn't fit.
The total length field of WAL record, xl_tot_len, is moved to the beginning
of the WAL record header, so that it is still always found on the first page
where a WAL record begins.
Bump WAL version number again as this is an incompatible change.
The continuation record only contained one field, xl_rem_len, so it makes
things simpler to just include it in the WAL page header. This wastes four
bytes on pages that don't begin with a continuation from previos page, plus
four bytes on every page, because of padding.
The motivation of this is to make it easier to calculate how much space a
WAL record needs. Before this patch, it depended on how many page boundaries
the record crosses. The motivation of that, in turn, is to separate the
allocation of space in the WAL from the copying of the record data to the
allocated space. Keeping the calculation of space required simple helps to
keep the critical section of allocating the space from WAL short. But that's
not included in this patch yet.
Bump WAL version number again, as this is an incompatible change.
The comments claimed that wasting the last segment made it easier to do
calculations with XLogRecPtrs, because you don't have problems representing
last-byte-position-plus-1 that way. In my experience, however, it only made
things more complicated, because the there was two ways to represent the
boundary at the beginning of a logical log file: logid = n+1 and xrecoff = 0,
or as xlogid = n and xrecoff = 4GB - XLOG_SEG_SIZE. Some functions were
picky about which representation was used.
Also, use a 64-bit segment number instead of the log/seg combination, to
point to a certain WAL segment. We assume that all platforms have a working
64-bit integer type nowadays.
This is an incompatible change in WAL format, so bumping WAL version number.
When I implemented the ginbuildempty() function as part of
implementing unlogged tables, I falsified the note in the header
comment for log_newpage. Although we could fix that up by changing
the comment, it seems cleaner to add a new function which is
specifically intended to handle this case. So do that.
We used to mimic the way a stack is constructed when descending the tree
during normal GiST inserts, but that was quite complicated during a buffered
build. It was also wrong: in GiST, the left-to-right relationships on
different levels might not match each other, so that when you know the
parent of a child page, you won't necessarily find the parent of the page to
the right of the child page by following the rightlinks at the parent level.
This sometimes led to "could not re-find parent" errors while building a
GiST index.
We now use a simple hash table to track the parent of every internal page.
Whenever a page is split, and downlinks are moved from one page to another,
we update the hash table accordingly. This is also better for performance
than the old method, as we never need to move right to re-find the parent
page, which could take a significant amount of time for buffers that were
created much earlier in the index build.
When we create a temporary copy of the old node buffer, in stack, we mustn't
leak that into any of the long-lived data structures. Before this patch,
when we called gistPopItupFromNodeBuffer(), it got added to the array of
"loaded buffers". After gistRelocateBuildBuffersOnSplit() exits, the
pointer added to the loaded buffers array points to garbage. Often that goes
unnotied, because when we go through the array of loaded buffers to unload
them, buffers with a NULL pageBuffer are ignored, which can often happen by
accident even if the pointer points to garbage.
This patch fixes that by marking the temporary copy in stack explicitly as
temporary, and refrain from adding buffers marked as temporary to the array
of loaded buffers.
While we're at it, initialize nodeBuffer->pageBlocknum to InvalidBlockNumber
and improve comments a bit. This isn't strictly necessary, but makes
debugging easier.
Users of asynchronous-commit mode expect there to be a guaranteed maximum
delay before an async commit's WAL records get flushed to disk. The
original version of the walwriter hibernation patch broke that. Add an
extra shared-memory flag to allow async commits to kick the walwriter out
of hibernation mode, without adding any noticeable overhead in cases where
no action is needed.
This patch modifies the walwriter process so that, when it has not found
anything useful to do for many consecutive wakeup cycles, it extends its
sleep time to reduce the server's idle power consumption. It reverts to
normal as soon as it's done any successful flushes. It's still true that
during any async commit, backends check for completed, unflushed pages of
WAL and signal the walwriter if there are any; so that in practice the
walwriter can get awakened and returned to normal operation sooner than the
sleep time might suggest.
Also, improve the checkpointer so that it uses a latch and a computed delay
time to not wake up at all except when it has something to do, replacing a
previous hardcoded 0.5 sec wakeup cycle. This also is primarily useful for
reducing the server's power consumption when idle.
In passing, get rid of the dedicated latch for signaling the walwriter in
favor of using its procLatch, since that comports better with possible
generic signal handlers using that latch. Also, fix a pre-existing bug
with failure to save/restore errno in walwriter's signal handlers.
Peter Geoghegan, somewhat simplified by Tom
The alternative of disallowing index-only scans in HS operation was
discussed, but the consensus was that it was better to treat marking
a page all-visible as a recovery conflict for snapshots that could still
fail to see XIDs on that page. We may in the future try to soften this,
so that we simply force index scans to do heap fetches in cases where
this may be an issue, rather than throwing a hard conflict.
The original syntax wasn't universally loved, and it didn't allow its
usage in CREATE TABLE, only ALTER TABLE. It now works everywhere, and
it also allows using ALTER TABLE ONLY to add an uninherited CHECK
constraint, per discussion.
The pg_constraint column has accordingly been renamed connoinherit.
This commit partly reverts some of the changes in
61d81bd28dbec65a6b144e0cd3d0bfe25913c3ac, particularly some pg_dump and
psql bits, because now pg_get_constraintdef includes the necessary NO
INHERIT within the constraint definition.
Author: Nikhil Sontakke
Some tweaks by me
It used to be case that lazy vacuum could call this function with only
a shared lock on the buffer, but neither lazy vacuum nor any other
code path does that any more. Simplify the code accordingly and clean
up some related, obsolete comments.
This patch fixes the other major compatibility-breaking limitation of
SPGiST, that it didn't store anything for null values of the indexed
column, and so could not support whole-index scans or "x IS NULL"
tests. The approach is to create a wholly separate search tree for
the null entries, and use fixed "allTheSame" insertion and search
rules when processing this tree, instead of calling the index opclass
methods. This way the opclass methods do not need to worry about
dealing with nulls.
Catversion bump is for pg_am updates as well as the change in on-disk
format of SPGiST indexes; there are some tweaks in SPGiST WAL records
as well.
Heavily rewritten version of a patch by Oleg Bartunov and Teodor Sigaev.
(The original also stored nulls separately, but it reused GIN code to do
so; which required undesirable compromises in the on-disk format, and
would likely lead to bugs due to the GIN code being required to work in
two very different contexts.)
The original API definition was incapable of supporting whole-index scans
because there was no way to invoke leaf-value reconstruction without
checking any qual conditions. Also, it was inefficient for
multiple-qual-condition scans because value reconstruction got done over
again for each qual condition, and because other internal work in the
consistent functions likewise had to be done for each qual. To fix these
issues, pass the whole scankey array to the opclass consistent functions,
instead of only letting them see one item at a time. (Essentially, the
loop over scankey entries is now inside the consistent functions not
outside them. This makes the consistent functions a bit more complicated,
but not unreasonably so.)
In itself this commit does nothing except save a few cycles in
multiple-qual-condition index scans, since we can't support whole-index
scans on SPGiST indexes until nulls are included in the index. However,
I consider this a must-fix for 9.2 because once we release it will get
very much harder to change the opclass API definition.
Comparing two xlog locations are useful for example when calculating
replication lag.
Euler Taveira de Oliveira, reviewed by Fujii Masao, and some cleanups
from me
Base backup follows recommended procedure, plus goes to great
lengths to ensure that partial page writes are avoided.
Jun Ishizuka and Fujii Masao, with minor modifications
Replication occurs only to memory on standby, not to disk,
so provides additional performance if user wishes to
reduce durability level slightly. Adds concept of multiple
independent sync rep queues.
Fujii Masao and Simon Riggs
