JSON_TABLE() allows JSON data to be converted into a relational view
and thus used, for example, in a FROM clause, like other tabular
data. Data to show in the view is selected from a source JSON object
using a JSON path expression to get a sequence of JSON objects that's
called a "row pattern", which becomes the source to compute the
SQL/JSON values that populate the view's output columns. Column
values themselves are computed using JSON path expressions applied to
each of the JSON objects comprising the "row pattern", for which the
SQL/JSON query functions added in 6185c9737c are used.
To implement JSON_TABLE() as a table function, this augments the
TableFunc and TableFuncScanState nodes that are currently used to
support XMLTABLE() with some JSON_TABLE()-specific fields.
Note that the JSON_TABLE() spec includes NESTED COLUMNS and PLAN
clauses, which are required to provide more flexibility to extract
data out of nested JSON objects, but they are not implemented here
to keep this commit of manageable size.
Author: Nikita Glukhov <n.gluhov@postgrespro.ru>
Author: Teodor Sigaev <teodor@sigaev.ru>
Author: Oleg Bartunov <obartunov@gmail.com>
Author: Alexander Korotkov <aekorotkov@gmail.com>
Author: Andrew Dunstan <andrew@dunslane.net>
Author: Amit Langote <amitlangote09@gmail.com>
Author: Jian He <jian.universality@gmail.com>
Reviewers have included (in no particular order):
Andres Freund, Alexander Korotkov, Pavel Stehule, Andrew Alsup,
Erik Rijkers, Zihong Yu, Himanshu Upadhyaya, Daniel Gustafsson,
Justin Pryzby, Álvaro Herrera, Jian He
Discussion: https://postgr.es/m/cd0bb935-0158-78a7-08b5-904886deac4b@postgrespro.ru
Discussion: https://postgr.es/m/20220616233130.rparivafipt6doj3@alap3.anarazel.de
Discussion: https://postgr.es/m/abd9b83b-aa66-f230-3d6d-734817f0995d%40postgresql.org
Discussion: https://postgr.es/m/CA+HiwqE4XTdfb1nW=Ojoy_tQSRhYt-q_kb6i5d4xcKyrLC1Nbg@mail.gmail.com
This changes the three callers to json_parse_manifest() to use
json_parse_manifest_incremental_chunk() if appropriate. In the case of
the backend caller, since we don't know the size of the manifest in
advance we always call the incremental parser.
Author: Andrew Dunstan
Reviewed-By: Jacob Champion
Discussion: https://postgr.es/m/7b0a51d6-0d9d-7366-3a1a-f74397a02f55@dunslane.net
This adds the infrastructure for using the new non-recursive JSON parser
in processing manifests. It's important that callers make sure that the
last piece of json handed to the incremental manifest parser contains
the entire last few lines of the manifest, including the checksum.
Author: Andrew Dunstan
Reviewed-By: Jacob Champion
Discussion: https://postgr.es/m/7b0a51d6-0d9d-7366-3a1a-f74397a02f55@dunslane.net
This parser uses an explicit prediction stack, unlike the present
recursive descent parser where the parser state is represented on the
call stack. This difference makes the new parser suitable for use in
incremental parsing of huge JSON documents that cannot be conveniently
handled piece-wise by the recursive descent parser. One potential use
for this will be in parsing large backup manifests associated with
incremental backups.
Because this parser is somewhat slower than the recursive descent
parser, it is not replacing that parser, but is an additional parser
available to callers.
For testing purposes, if the build is done with -DFORCE_JSON_PSTACK, all
JSON parsing is done with the non-recursive parser, in which case only
trivial regression differences in error messages should be observed.
Author: Andrew Dunstan
Reviewed-By: Jacob Champion
Discussion: https://postgr.es/m/7b0a51d6-0d9d-7366-3a1a-f74397a02f55@dunslane.net
Commit 619bc23a1a introduced
WARNING: Project targets '>=0.54' but uses feature introduced in '0.55.0': Passing executable/found program object to script parameter of add_dist_script
Work around that by wrapping the offending line in a meson version check.
Author: Tristan Partin <tristan@neon.tech>
Discussion: https://www.postgresql.org/message-id/flat/D096Q3NFFVH1.1T5RE4MOO9ZFH%40neon.tech
This would cause the timestamp values used by emit_log_hook and all the
other log destinations to differ, because the timestamps are reset
before sending the logs to the server and after calling the hook.
This change matters for emit_log_hook when generating log information
with 'n' or 'm' in log_line_prefix through log_status_format(), or when
doing direct calls to get_formatted_log_time() like in the JSON or CSV
log formats.
While on it, this commit fixes a couple of comments related to the
formatted timestamps where the JSON was not mentioned. Oversight in
dc686681e0, that I have noticed while reviewing this patch.
Author: Kambam Vinay, Michael Paquier
Discussion: https://postgr.es/m/CANiRfmsK36A0i8mnQtzaxhSm3CUCimPwJPp4WQNq53OdSNkgWg@mail.gmail.com
To allow the use of the read stream API added in b5a9b18cd for
sequential scans on heap tables, here we make some adjustments to make
that change less invasive and perhaps make the code easier to follow in
the process.
Here heapgetpage() gets broken into two functions:
1) The part which reads the block has now been moved into a function
named heapfetchbuf().
2) The part which performed pruning and populated the scan's
rs_vistuples[] array is now moved into a new function named
heap_prepare_pagescan().
The functionality provided by heap_prepare_pagescan() was only ever
required by SO_ALLOW_PAGEMODE scans, so the branching that was
previously done in heapgetpage() is no longer needed as we simply just
don't call heap_prepare_pagescan() from heapgettup() in the refactored
code.
Author: Melanie Plageman
Discussion: https://postgr.es/m/CAAKRu_YtXJiYKQvb5JsA2SkwrsizYLugs4sSOZh3EAjKUg=gEQ@mail.gmail.com
emit_tap_output_v() includes some fprintf() calls for some output
related to the TAP protocol, that may clobber errno and break %m. This
commit makes the logging of pg_regress smarter by saving errno before
restoring it in vfprintf() where the input strings are used, removing
the need for strerror(). All logs are switched to %m rather than
strerror(), shaving some code.
This was not a problem until now as pg_regress.c did not use %m, but the
change is simple enough that we have no reason to not support this
placeholder, and that will avoid future mistakes if new logs that
include %m are added.
Author: Dagfinn Ilmari Mannsåker
Reviewed-by: Peter Eisentraunt, Michael Paquier
Discussion: https://postgr.es/m/87sf13jhuw.fsf@wibble.ilmari.org
EXPLAIN (ANALYZE, SERIALIZE) allows collection of statistics about
the volume of data emitted by a query, as well as the time taken
to convert the data to the on-the-wire format. Previously there
was no way to investigate this without actually sending the data
to the client, in which case network transmission costs might
swamp what you wanted to see. In particular this feature allows
investigating the costs of de-TOASTing compressed or out-of-line
data during formatting.
Stepan Rutz and Matthias van de Meent,
reviewed by Tomas Vondra and myself
Discussion: https://postgr.es/m/ca0adb0e-fa4e-c37e-1cd7-91170b18cae1@gmx.de
After this change we have XLogCtl->logWriteResult and ->logFlushResult.
There's no functional change, other than the fact that the assignment
from shared memory to local is no longer done via struct assignment, but
instead using a macro that copies each member separately.
The current representation is inconvenient going forward; notably, we
would like to add a new member "Copy" (to keep track of the last
position copied into WAL buffers), so the symmetry between the values in
shared memory vs. those in local would be lost.
This also gives us freedom to later change the concurrency model for the
values in shared memory: we can make them use atomics instead of relying
on the info_lck spinlock.
Reviewed-by: Bharath Rupireddy <bharath.rupireddyforpostgres@gmail.com>
Discussion: https://postgr.es/m/202404031119.cd2kugjk2vho@alvherre.pgsql
If there aren't many bytes to process, the function call overhead
of the optimized implementation isn't worth taking, so instead we
inline a loop that consults pg_number_of_ones in that case. If
there are many bytes to process, we accept the function call
overhead because the optimized versions are likely to be faster.
The threshold at which we use the optimized implementation is set
to the smallest amount of data required to use special popcount
instructions.
Reviewed-by: Alvaro Herrera, Tom Lane
Discussion: https://postgr.es/m/20240402155301.GA2750455%40nathanxps13
Execute both freezing and pruning of tuples in the same
heap_page_prune() function, now called heap_page_prune_and_freeze(),
and emit a single WAL record containing all changes. That reduces the
overall amount of WAL generated.
This moves the freezing logic from vacuumlazy.c to the
heap_page_prune_and_freeze() function. The main difference in the
coding is that in vacuumlazy.c, we looked at the tuples after the
pruning had already happened, but in heap_page_prune_and_freeze() we
operate on the tuples before pruning. The heap_prepare_freeze_tuple()
function is now invoked after we have determined that a tuple is not
going to be pruned away.
VACUUM no longer needs to loop through the items on the page after
pruning. heap_page_prune_and_freeze() does all the work. It now
returns the list of dead offsets, including existing LP_DEAD items, to
the caller. Similarly it's now responsible for tracking 'all_visible',
'all_frozen', and 'hastup' on the caller's behalf.
Author: Melanie Plageman <melanieplageman@gmail.com>
Discussion: https://www.postgresql.org/message-id/20240330055710.kqg6ii2cdojsxgje@liskov
In preparation of freezing and counting tuples which are not
candidates for pruning, split heap_prune_record_unchanged() into
multiple functions, depending the kind of line pointer. That's not too
interesting right now, but makes the next commit smaller.
Recording the lowest soon-to-be prunable xid is one of the actions we
take for unchanged LP_NORMAL item pointers but not for others, so move
that to the new heap_prune_record_unchanged_lp_normal() function. The
next commit will add more actions to these functions.
Author: Melanie Plageman <melanieplageman@gmail.com>
Discussion: https://www.postgresql.org/message-id/20240330055710.kqg6ii2cdojsxgje@liskov
Bug in commit 53c2a97a9266: we failed to acquire the correct SLRU bank
lock when iterating to zero-out intermediate pages in predicate.c.
Rewrite the code block so that we follow the locking protocol correctly.
Also update an outdated comment in the same file -- SerialSLRULock
exists no more.
Reported-by: Alexander Lakhin <exclusion@gmail.com>
Reviewed-by: Dilip Kumar <dilipbalaut@gmail.com>
Discussion: https://postgr.es/m/2a25eaf4-a3a4-5fd1-6241-9d7c73142085@gmail.com
06c418e163 introduced pg_wal_replay_wait() procedure allowing to wait for
the particular LSN to be replayed on standby. The waiters were stored in
the flat array. Even though scanning small arrays is fast, that might be a
problem at scale (a lot of waiting processes).
This commit replaces the flat shared memory array with the pairing heap,
which holds the waiter with the least LSN at the top. This gives us O(log N)
complexity for both inserting and removing waiters.
Reported-by: Alvaro Herrera
Discussion: https://postgr.es/m/202404030658.hhj3vfxeyhft%40alvherre.pgsql
Project policy is to not leave global objects behind after a regress
test run. This was found as a result of the development of a patch
to make pg_regress detect such leftovers automatically, which in the
end was withdrawn due to issues with parallel runs.
Discussion: https://postgr.es/m/E1phvk7-000VAH-7k@gemulon.postgresql.org
We were directly copying the LSN locations while syncing the slots on the
standby. Now, it is possible that at some particular restart_lsn there are
some running xacts, which means if we start reading the WAL from that
location after promotion, we won't reach a consistent snapshot state at
that point. However, on the primary, we would have already been in a
consistent snapshot state at that restart_lsn so we would have just
serialized the existing snapshot.
To avoid this problem we will use the advance_slot functionality unless
the snapshot already exists at the synced restart_lsn location. This will
help us to ensure that snapbuilder/slot statuses are updated properly
without generating any changes. Note that the synced slot will remain as
RS_TEMPORARY till the decoding from corresponding restart_lsn can reach a
consistent snapshot state after which they will be marked as
RS_PERSISTENT.
Per buildfarm
Author: Hou Zhijie
Reviewed-by: Bertrand Drouvot, Shveta Malik, Bharath Rupireddy, Amit Kapila
Discussion: https://postgr.es/m/OS0PR01MB5716B3942AE49F3F725ACA92943B2@OS0PR01MB5716.jpnprd01.prod.outlook.com
This adds errcodes to a set of PANIC and FATAL errors in xlog.c
and relcache.c, which previously had no errcode at all set, in
order to make fleetwide analysis of errorlogs easier. There are
many more ereport/elogs left which could benefit from having an
errcode but this at least makes a dent in the issue.
Author: Nazir Bilal Yavuz <byavuz81@gmail.com>
Discussion: https://postgr.es/m/CAN55FZ1k8LgLEqncPGmz_fWnrobV6bjABOTH4tOWta6xNcPQig@mail.gmail.com
Presently, the archiver process restarts when an archive callback
ERRORs. To avoid this, archive module authors can use sigsetjmp(),
manage a memory context, etc., but that requires a lot of extra
code that will likely look roughly the same between modules. This
commit adds basic archive callback ERROR handling to pgarch.c so
that module authors won't ordinarily need to worry about this.
While this built-in handler attempts to clean up anything that an
archive module could conceivably have left behind, it is possible
that some modules are doing unexpected things that require
additional cleanup. Module authors should be sure to do any extra
required cleanup in a PG_CATCH block within the archiving callback.
The archiving callback is now called in a short-lived memory
context that the archiver process resets between invocations. If a
module requires longer-lived storage, it must maintain its own
memory context.
Thanks to these changes, the basic_archive module can be greatly
simplified.
Suggested-by: Andres Freund
Reviewed-by: Andres Freund, Yong Li
Discussion: https://postgr.es/m/20230217215624.GA3131134%40nathanxps13
Previously, when selecting the transaction to evict during logical
decoding, we check all transactions to find the largest
transaction. This could lead to a significant replication lag
especially in the case where there are many subtransactions.
This commit improves the eviction algorithm in ReorderBuffer using the
max-heap with transaction size as the key to efficiently find the
largest transaction.
The max-heap starts with empty. While the max-heap is empty, we don't
do anything for the max-heap when updating the memory
counter. Therefore, we get the largest transaction in O(N) time, where
N is the number of transactions including top-level transactions and
subtransactions.
We build the max-heap just before selecting the largest transactions
if the number of transactions being decoded is higher than the
threshold, MAX_HEAP_TXN_COUNT_THRESHOLD. After building the max-heap,
we also update the max-heap when updating the memory counter. The
intention is to efficiently find the largest transaction in O(1) time
instead of incurring the cost of memory counter updates (O(log
N)). Once the number of transactions got lower than the threshold, we
reset the max-heap.
The performance benchmark results showed significant speed up (more
than x30 speed up on my machine) in decoding a transaction with 100k
subtransactions, whereas there is no visible overhead in other cases.
Reviewed-by: Amit Kapila, Hayato Kuroda, Vignesh C, Ajin Cherian,
Tomas Vondra, Shubham Khanna, Peter Smith, Álvaro Herrera,
Euler Taveira
Discussion: https://postgr.es/m/CAD21AoAfKTgrBrLq96GcTv9d6k97zaQcDM-rxfKEt4GSe0qnaQ%40mail.gmail.com
Previously, binaryheap didn't support updating a key and removing a
node in an efficient way. For example, in order to remove a node from
the binaryheap, the caller had to pass the node's position within the
array that the binaryheap internally has. Removing a node from the
binaryheap is done in O(log n) but searching for the key's position is
done in O(n).
This commit adds a hash table to binaryheap in order to track the
position of each nodes in the binaryheap. That way, by using newly
added functions such as binaryheap_update_up() etc., both updating a
key and removing a node can be done in O(1) on an average and O(log n)
in worst case. This is known as the indexed binary heap. The caller
can specify to use the indexed binaryheap by passing indexed = true.
The current code does not use the new indexing logic, but it will be
used by an upcoming patch.
Reviewed-by: Vignesh C, Peter Smith, Hayato Kuroda, Ajin Cherian,
Tomas Vondra, Shubham Khanna
Discussion: https://postgr.es/m/CAD21AoDffo37RC-eUuyHJKVEr017V2YYDLyn1xF_00ofptWbkg%40mail.gmail.com
Since 66c0185a3, the planner is able to use Merge Append -> Unique to
implement UNION queries and each subquery is prompted to produce Paths
correctly sorted by the UNION's targetlist.
Here we remove some now redundant code which was zeroing the
tuple_fraction at the parent level. This will allow the planner to
consider cheap startup paths when planning the UNION's subqueries.
EXCEPT and INTERSECT set operations still have the tuple_fraction zeroed
in generate_nonunion_paths(). These operations currently always read
all of their subqueries' tuples.
Reported-by: Tom Lane
Discussion: https://postgr.es/m/3703023.1711654574@sss.pgh.pa.us
pg_wal_replay_wait() is to be used on standby and specifies waiting for
the specific WAL location to be replayed before starting the transaction.
This option is useful when the user makes some data changes on primary and
needs a guarantee to see these changes on standby.
The queue of waiters is stored in the shared memory array sorted by LSN.
During replay of WAL waiters whose LSNs are already replayed are deleted from
the shared memory array and woken up by setting of their latches.
pg_wal_replay_wait() needs to wait without any snapshot held. Otherwise,
the snapshot could prevent the replay of WAL records implying a kind of
self-deadlock. This is why it is only possible to implement
pg_wal_replay_wait() as a procedure working in a non-atomic context,
not a function.
Catversion is bumped.
Discussion: https://postgr.es/m/eb12f9b03851bb2583adab5df9579b4b%40postgrespro.ru
Author: Kartyshov Ivan, Alexander Korotkov
Reviewed-by: Michael Paquier, Peter Eisentraut, Dilip Kumar, Amit Kapila
Reviewed-by: Alexander Lakhin, Bharath Rupireddy, Euler Taveira
If temp tables have dependencies (such as sequences) then it's
possible for autovacuum's cleanup of orphan temp tables to deadlock
against an incoming backend that's trying to clean out the temp
namespace for its own use. That can happen because RemoveTempRelations'
performDeletion call can visit objects within the namespace in
an order different from the order in which a per-table deletion
will visit them.
To fix, observe that performDeletion will begin by taking an exclusive
lock on the temp namespace (even though it won't actually delete it).
So, if we can get a shared lock on the namespace, we can be sure we're
not running concurrently with RemoveTempRelations, while also not
conflicting with ordinary use of the namespace. This requires
introducing a conditional version of LockDatabaseObject, but that's no
big deal. (It's surprising we've got along without that this long.)
Report and patch by Mikhail Zhilin. Back-patch to all supported
branches.
Discussion: https://postgr.es/m/c43ce028-2bc2-4865-9b89-3f706246eed5@postgrespro.ru
Make this code a little less confusing by defining a separate macro
that controls whether we'll use ResourceOwner facilities to track
the existence of a pg_hmac_ctx context.
The proximate reason to touch this is that since b8bff07da, we got
"unused function" warnings if building with older OpenSSL, because
the #if guards around the ResourceOwner wrapper function definitions
were different from those around the calls of those functions.
Pulling the ResourceOwner machinations outside of the #ifdef HAVE_xxx
guards fixes that and makes the code clearer too.
Discussion: https://postgr.es/m/1394271.1712016101@sss.pgh.pa.us
After installing the SIGINT handler in psql, SIGINT can no longer cancel
database reconnections. For instance, if the user starts a reconnection
and then needs to do some form of interaction (ie psql is polling),
there is no way to cancel the reconnection process currently.
Use PQconnectStartParams() in order to insert a cancel_pressed check
into the polling loop.
Tristan Partin, reviewed by Gurjeet Singh, Heikki Linnakangas, Jelte
Fennema-Nio, and me.
Discussion: http://postgr.es/m/D08WWCPVHKHN.3QELIKZJ2D9RZ@neon.tech
This is useful when connecting to a database asynchronously via
PQconnectStart(), since it handles deciding between poll() and
select(), and some of the required boilerplate.
Tristan Partin, reviewed by Gurjeet Singh, Heikki Linnakangas, Jelte
Fennema-Nio, and me.
Discussion: http://postgr.es/m/D08WWCPVHKHN.3QELIKZJ2D9RZ@neon.tech
Introduce an abstraction allowing relation data to be accessed as a
stream of buffers, with an implementation that is more efficient than
the equivalent sequence of ReadBuffer() calls.
Client code supplies a callback that can say which block number it wants
next, and then consumes individual buffers one at a time from the
stream. This division puts read_stream.c in control of how far ahead it
can see and allows it to read clusters of neighboring blocks with
StartReadBuffers(). It also issues POSIX_FADV_WILLNEED advice ahead of
time when random access is detected.
Other variants of I/O stream will be proposed in future work (for
example to support recovery, whose LsnReadQueue device in
xlogprefetcher.c is a distant cousin of this code and should eventually
be replaced by this), but this basic API is sufficient for many common
executor usage patterns involving predictable access to a single fork of
a single relation.
Several patches using this API are proposed separately.
This stream concept is loosely based on ideas from Andres Freund on how
we should pave the way for later work on asynchronous I/O.
Author: Thomas Munro <thomas.munro@gmail.com>
Author: Heikki Linnakangas <hlinnaka@iki.fi> (contributions)
Author: Melanie Plageman <melanieplageman@gmail.com> (contributions)
Suggested-by: Andres Freund <andres@anarazel.de>
Reviewed-by: Heikki Linnakangas <hlinnaka@iki.fi>
Reviewed-by: Melanie Plageman <melanieplageman@gmail.com>
Reviewed-by: Nazir Bilal Yavuz <byavuz81@gmail.com>
Reviewed-by: Andres Freund <andres@anarazel.de>
Tested-by: Tomas Vondra <tomas.vondra@enterprisedb.com>
Discussion: https://postgr.es/m/CA+hUKGJkOiOCa+mag4BF+zHo7qo=o9CFheB8=g6uT5TUm2gkvA@mail.gmail.com
Break ReadBuffer() up into two steps. StartReadBuffers() and
WaitReadBuffers() give us two main advantages:
1. Multiple consecutive blocks can be read with one system call.
2. Advice (hints of future reads) can optionally be issued to the
kernel ahead of time.
The traditional ReadBuffer() function is now implemented in terms of
those functions, to avoid duplication.
A new GUC io_combine_limit is defined, and the functions for limiting
per-backend pin counts are made into public APIs. Those are provided
for use by callers of StartReadBuffers(), when deciding how many buffers
to read at once. The following commit will add a higher level mechanism
for doing that automatically with a practical interface.
With some more infrastructure in later work, StartReadBuffers() could
be extended to start real asynchronous I/O instead of just issuing
advice and leaving WaitReadBuffers() to do the work synchronously.
Author: Thomas Munro <thomas.munro@gmail.com>
Author: Andres Freund <andres@anarazel.de> (some optimization tweaks)
Reviewed-by: Melanie Plageman <melanieplageman@gmail.com>
Reviewed-by: Heikki Linnakangas <hlinnaka@iki.fi>
Reviewed-by: Nazir Bilal Yavuz <byavuz81@gmail.com>
Reviewed-by: Dilip Kumar <dilipbalaut@gmail.com>
Reviewed-by: Andres Freund <andres@anarazel.de>
Tested-by: Tomas Vondra <tomas.vondra@enterprisedb.com>
Discussion: https://postgr.es/m/CA+hUKGJkOiOCa+mag4BF+zHo7qo=o9CFheB8=g6uT5TUm2gkvA@mail.gmail.com
Previously, we used a simple array for storing dead tuple IDs during
lazy vacuum, which had a number of problems:
* The array used a single allocation and so was limited to 1GB.
* The allocation was pessimistically sized according to table size.
* Lookup with binary search was slow because of poor CPU cache and
branch prediction behavior.
This commit replaces that array with the TID store from commit
30e144287a.
Since the backing radix tree makes small allocations as needed, the
1GB limit is now gone. Further, the total memory used is now often
smaller by an order of magnitude or more, depending on the
distribution of blocks and offsets. These two features should make
multiple rounds of heap scanning and index cleanup an extremely rare
event. TID lookup during index cleanup is also several times faster,
even more so when index order is correlated with heap tuple order.
Since there is no longer a predictable relationship between the number
of dead tuples vacuumed and the space taken up by their TIDs, the
number of tuples no longer provides any meaningful insights for users,
nor is the maximum number predictable. For that reason this commit
also changes to byte-based progress reporting, with the relevant
columns of pg_stat_progress_vacuum renamed accordingly to
max_dead_tuple_bytes and dead_tuple_bytes.
For parallel vacuum, both the TID store and supplemental information
specific to vacuum are shared among the parallel vacuum workers. As
with the previous array, we don't take any locks on TidStore during
parallel vacuum since writes are still only done by the leader
process.
Bump catalog version.
Reviewed-by: John Naylor, (in an earlier version) Dilip Kumar
Discussion: https://postgr.es/m/CAD21AoAfOZvmfR0j8VmZorZjL7RhTiQdVttNuC4W-Shdc2a-AA%40mail.gmail.com
66c0185a3 adjusted the UNION planner to request that union child queries
produce Paths correctly ordered to implement the UNION by way of
MergeAppend followed by Unique. The code there made a bad assumption
that if the root->parent_root->parse had setOperations set that the
query must be the child subquery of a set operation. That's not true
when it comes to planning a non-inlined CTE which is parented by a set
operation. This causes issues as the CTE's targetlist has no
requirement to match up to the SetOperationStmt's groupClauses
Fix this by adding a new parameter to both subquery_planner() and
grouping_planner() to explicitly pass the SetOperationStmt only when
planning set operation child subqueries.
Thank you to Tom Lane for helping to rationalize the decision on the
best function signature for subquery_planner().
Reported-by: Alexander Lakhin
Discussion: https://postgr.es/m/242fc7c6-a8aa-2daf-ac4c-0a231e2619c1@gmail.com
Andrew Dunstan