Tag Archives: performance

This post is about one of the fundamentally important properties of a database: how IO is done. The test case I studied is doing a simple full table scan of a single large table. In both Oracle and postgres the table doesn’t have any indexes or constraints, which is not a realistic example, but this doesn’t change the principal topic of the study: doing a table scan.

I used a publicly available dataset from the US bureau of transportation statistics called FAF4.5.1_database.zip
The zipped file is 347MB, unzipped size 1.7GB.

In both cases Oracle Linux 7.7 (64 bit) is used, running in VirtualBox, with the storage being a USB3 SSD. Number of CPUs is 4, memory size is 6G. Filesystem type: xfs.
The Oracle version used is Oracle 19.5, the Postgresql version used is 12.1.
For Postgresql, the postgresql.conf file is not changed, except for max_parallel_workers_per_gather which is set to 0 to make postgres use a single process.
For Oracle, the parameters that I think are important: filesystemio_options=’setall’. Oracle is used filesystem based (so no ASM).

This is the table definition for Oracle:

create table faf451 (
  fr_origin varchar2(3),
  dms_orig varchar2(3),
  dms_dest varchar2(3),
  fr_dest varchar2(3),
  fr_inmode varchar2(1),
  dms_mode varchar2(1),
  fr_outmode varchar2(1),
  sctg2 varchar2(2),
  trade_type varchar2(1),
  tons number,
  value number,
  tmiles number,
  curval number,
  wgt_dist number,
  year varchar2(4)
);

This is the table definition for Postgresql:

create table faf451 (
  fr_origin varchar(3),
  dms_orig varchar(3),
  dms_dest varchar(3),
  fr_dest varchar(3),
  fr_inmode varchar(1),
  dms_mode varchar(1),
  fr_outmode varchar(1),
  sctg2 varchar(2),
  trade_type varchar(1),
  tons double precision,
  value double precision,
  tmiles double precision,
  curval double precision,
  wgt_dist double precision,
  year varchar(4)
);

In order for the data to be easy loadable into postgres using copy from, I had to remove ‘””‘ (double double quotes) for the empty numeric fields. In oracle I could say “optionally enclosed by ‘”‘”. For Oracle I used an external table definition to load the data.

Now, before doing any benchmarks, I have an idea where this is going. Oracle is using direct IO (DIO) so linux page cache management and “double buffering” are avoided. Also, oracle will be doing asynchronous IO (AIO), which means submitting is separated from waiting for the notification that the submitted IOs are ready, and on top of that oracle will submit multiple IO requests at the same time. And again on top of that, oracle does multi-block IO, which means that instead of requesting each 8K database block individually, it will group adjacent blocks and request for these in one go, up to a size of combined blocks of 1MB, which means it can requests up to 128 8K blocks in one IO. Postgres will request every block synchronous, so 1 8K block at a time, and waiting for each request to finish. That makes me have a strong idea where this is going.

It should be noted that postgres explicitly is depending on the operating system page cache for buffering as a design principle. Because of DIO, blocks that are read by oracle are not cached in the operating system page cache.

I executed my benchmark in the following way:
– A run for every size is executed 5 times.
– At the start of every run for a certain size (so before every “batch” of 5 runs), the page cache is flushed: (echo 3 > /proc/sys/vm/drop_caches).
– Before each individual run, the database cache is flushed (systemctl restart postgresql-12 for postgres, alter system flush buffer_cache for oracle).

I started off with 2G from the dataset, and then simply performed a ‘copy from’ again to load the same dataset into the table in postgres. Oracle required a bit more of work. Oracle was able to save the same data in way less blocks; the size became 1.18G. In order to have both postgres and oracle scan the same amount of data, I calculated roughly how much rows I needed to add to the table to make it 2G, and copied that table to save it as a 2G table, so I could insert that table to increase the size of the test table by 2G. This way in both oracle and postgres I could test with a 2G table and add 2G at a time until I reached 20G.

These are the results. As you can see in the legenda: oracle is orange, postgres is blue.
(click graph to load full picture)

What we see, is that postgres is a bit slower with the first run of 5 for the smaller dataset sizes, which becomes less visible with larger datasets.
Also, postgres is way faster if the dataset fits into the page cache and it has been read into it. This is logical because postgres explicitly uses the page cache as a secondary cache, and the test is the only activity on this server, so it hasn’t been flushed by other activity.

What was totally shocking to me, is postgres is performing alike oracle and both roughly are able to perform at the maximum IO speed of my disk: 300MB/s, especially when the dataset is bigger, alias beyond the page cache size.

It wasn’t shocking that oracle could reach the total bandwidth of the disk: oracle uses all the techniques to optimise IO for bandwidth. But how can postgres do the same, NOT deploying these techniques, reading 8K at a time??

The first thing to check is whether postgres is doing something else than I suspected. This can simply be checked with strace:

poll_wait(3, [{EPOLLIN, {u32=18818136, u64=18818136}}], 1, -1) = 1
recvfrom(11, "Q\0\0\0!select count(*) from faf451"..., 8192, 0, NULL, NULL) = 34
lseek(20, 0, SEEK_END)                  = 335740928
lseek(20, 0, SEEK_END)                  = 335740928
kill(1518, SIGUSR1)                     = 0
pread64(5, "\f\0\0\0\310ILc\0\0\0\0D\1\210\1\0 \4 \0\0\0\0\230\237\312\0000\237\312\0"..., 8192, 846061568) = 8192
pread64(5, "\f\0\0\0HcLc\0\0\0\0D\1\210\1\0 \4 \0\0\0\0\230\237\312\0000\237\312\0"..., 8192, 846069760) = 8192
pread64(5, "\f\0\0\0\260|Lc\0\0\0\0D\1\210\1\0 \4 \0\0\0\0\230\237\312\0000\237\312\0"..., 8192, 846077952) = 8192
pread64(5, "\f\0\0\0000\226Lc\0\0\0\0D\1\210\1\0 \4 \0\0\0\0\230\237\312\0000\237\312\0"..., 8192, 846086144) = 8192
…etc…

The above strace output shows only 4 rows of pread64() calls, but this goes on. So no “secret” optimisation there.

Luckily, my VM has a new enough version of Linux for it to be able to use eBPF, so I can use biosnoop. Biosnoop is a tool to look at IO on one of the lower layers of the linux kernel, the block device interface (hence ‘bio’). This is the biosnoop output:

# /usr/share/bcc/tools/biosnoop
TIME(s)        COMM           PID    DISK    T  SECTOR    BYTES   LAT(ms)
0.000000000    postmaster     4143   sdb     R  66727776  708608     5.51
0.006419000    postmaster     4143   sdb     R  66731720  77824     11.06
0.006497000    postmaster     4143   sdb     R  66734432  786432    11.03
0.011550000    postmaster     4143   sdb     R  66731872  1310720   16.17
0.013470000    postmaster     4143   sdb     R  66729160  1310720   18.86
0.016439000    postmaster     4143   sdb     R  66735968  1310720   14.61
0.019220000    postmaster     4143   sdb     R  66738528  786432    15.20

Wow…so here it’s doing IOs of up to 1MB! So somewhere between postgres itself and the block device, the IOs magically grew to sizes up to 1MB…that’s weird. The only thing that sits between postgres and the block device is the linux kernel, which includes page cache management.

To get an insight into that, I ran ‘perf record -g -p PID’ during the scan, and then perf report to look at the recorded perf data. This is what is I found:

Samples: 21K of event 'cpu-clock', Event count (approx.): 5277000000
  Children      Self  Command     Shared Object       Symbol                                                                  ◆
-   41.84%     3.63%  postmaster  libpthread-2.17.so  [.] __pread_nocancel                                                    ▒
   - 38.20% __pread_nocancel                                                                                                  ▒
      - 38.08% entry_SYSCALL_64_after_hwframe                                                                                 ▒
         - 37.95% do_syscall_64                                                                                               ▒
            - 35.87% sys_pread64                                                                                              ▒
               - 35.51% vfs_read                                                                                              ▒
                  - 35.07% __vfs_read                                                                                         ▒
                     - 34.97% xfs_file_read_iter                                                                              ▒
                        - 34.69% __dta_xfs_file_buffered_aio_read_3293                                                        ▒
                           - 34.32% generic_file_read_iter                                                                    ▒
                              - 21.10% page_cache_async_readahead                                                             ▒
                                 - 21.04% ondemand_readahead                                                                  ▒
                                    - 20.99% __do_page_cache_readahead                                                        ▒
                                       + 14.14% __dta_xfs_vm_readpages_3179                                                   ▒
                                       + 5.07% __page_cache_alloc                                                             ▒
                                       + 0.97% radix_tree_lookup                                                              ▒
                                       + 0.54% blk_finish_plug                                                                ▒

If you look at rows 13-15 you see that the kernel is performing readahead. This is an automatic function in the linux kernel which looks if the requests are sequential of nature, and when that’s true performs readahead, so that the scan is made faster.

Where does the log writer spend its time on?

January 2, 2020

Exadata, Linux, Oracle, Oracle EE, Oracle XE

2 Comments

The Oracle database log writer is the process that fundamentally influences database change performance. Under normal circumstances the log writer must persist the changes made to the blocks before the actual change is committed. Therefore, it’s vitally important to understand what the log writer is exactly doing. This is widely known by the Oracle database community.

The traditional method for looking at log writer performance is looking at the wait event ‘log file parallel write’ and the CPU time, and comparing that to the ‘log file sync’ alias “commit wait time”. If ‘log file parallel write’ and ‘log file sync’ roughly match, a commit is waiting on the log writer IO latency, if it isn’t then it’s unclear, and things get vague.

Now we get to the gist of this blogpost: since Oracle 12, there are additional statistics in V$SESSTAT and v$SYSSTAT that describe where the log writer spending it’s time on. However, and sadly, these are not documented by Oracle. The statistics are:

– redo write gather time
– redo write schedule time
– redo write finish time
– redo write time (usec)
– redo write total time

Probably ‘redo write time (usec)’ doesn’t belong to these statistics because the timing is as good as the same as ‘redo write finish time’. Also the statistic ‘redo write time (usec)’ is updated by ‘kcrfw_postprocess_write’ right after the main writing function (kcrfw_do_write), whilst the other redo write timing statistics are updated/made visible by ksu_dispatch_tac outside of the actual writing functions, at the approximate rate of every 3 seconds.

All these statistics contain cumulative timing figure in microseconds (usec).

All these timings start at the start of the function kcrfw_redo_write_driver. This is the main function that handles all the redo writing, including investigating the fixed area structure (kcrfs_) and the shared pool structures (made visible using X$KCRFSTRAND) to see if there are changes.

1. redo write gather time.
This times reading the structures that control allocation of the redo structures, validating foreground processes are done writing the change vectors and applying these to the blocks, updating the LWN SCN. Part of this also happens when there is nothing to write, but then the time is not added to this statistic, only if it found something to write.

2. redo write schedule time.
This includes the time of 1, but now it found something to write, entered the writing function (kcrfw_do_write), and updated the ksfd IO related structures for writing.

3. redo write finish time.
This includes the time of 1 and 2, but this additionally times all the IO related functions for writing the changes it found in the public redo strands, as indicated by the shared pool and fixed area structures.

4. redo write time (usec).
This probably is not part of this group of statistics, but it works in exactly the same way: it’s a cumulative figure, it counts time since the start of the kcrfw_redo_write_driver and includes a tiny bit more code, it times until it has left the kcrfw_do_write function and enters the kcrfw_postprocess_write function.

5. redo write total time.
This includes the time of 1, 2 and 3, but additionally times the kcrfw_post function, which performs the signalling of the foreground processes that have committed and wait for a semop() signal.

This means that the Oracle database since Oracle 12 provides the timing information to see where the log writer is spending it’s time on. This can prevent a lot of guesswork.

Oracle wait event ‘log file parallel write’ change

December 31, 2019

Exadata, Linux, Oracle, Oracle EE, Oracle XE

Oracle wait event ‘TCP Socket (KGAS)’

October 21, 2018

Exadata, Linux, Networking, Oracle, Oracle EE, Oracle XE

Summary

The basic message of this article is not surprising, nor does it conflict with current knowledge. Whenever you see a wait event ‘TCP Socket (KGAS)’, it means a foreground process is performing TCP networking via PLSQL. This wait event is a single event for creating, sending, receiving and closing a connection.

The true information of this article is how you can use the p1 value of the event to learn what actually the foreground is doing, and thus should give you more information to troubleshoot in the case of long waiting times.

TCP Socket (KGAS) p1 values:
1 - ?
2 - perform DNS lookup and create connection
3 - close connection
4 - copy TCP response into process space
5 - send request
6 - wait for TCP response to become available
7 - ?
8 - ? \
9 - ? | According to documentation, windows only in a 'special thread'.
10- ? /

All about headroom and mandatory patching before June 2019

July 29, 2018

Exadata, Oracle, Oracle EE, Oracle XE

13 Comments

This post was triggered upon rereading a blogpost by Mike Dietrich called databases need patched minimum april 2019. Mike’s blogpost makes it clear this is about databases that are connected using database links, and that:
– Newer databases do not need additional patching for this issue (11.2.0.4, 12.1.0.2, 12.2 and newer).
– Recent PSU patches contain a fix for certain older versions (11.1.0.7, 11.2.0.3 and 12.1.0.1).
– This means versions 11.2.0.2 and earlier 11.2 versions, 11.1.0.6 and earlier and anything at version 10 or earlier can not be fixed and thus are affected.

But what is the actual issue?

The first link in the article: Recommended patches and actions for Oracle databases versions 12.1.0.1, 11.2.0.3 and earlier – before June 2019 (Doc ID 2361478.1) provides essentially the same information as Mike’s blogpost, however it additionaly mentions that the interoperability of database clients with database servers is not affected.

Mike’s article mentions the following:
The alert refers to an SCN issue which came up a while ago. The system change number (SCN) is a logical, internal timestamp used by the Oracle Database. SCNs order events that occur within the database, which is necessary to satisfy the ACID properties of a transaction. The database uses SCNs to query and track changes.

So I guess it has something to do with SCNs. Most of the links are about SCNs. The MOS article that is most descriptive is: System Change Number (SCN), Headroom, Security and Patch Information (Doc ID 1376995.1).

This article describes a lot of details:
– SCNs are necessary for the database to keep changes organised. I got an article that explains that SCNs are not unique to a transaction, but are “just” granular enough to keep changes organised.
– SCNs are an ever increasing number. SCNs are never decreased!
– SCNs have a hard limit, which version specific, and is based on the number of bits for the number:
– – The general limit is 2^48.
– – From version 12.2 on, with compatibility set to 12.2 or higher, the limit is 2^63.
– SCNs have a per second increasing soft limit, called ‘RSL’ or ‘reasonable SCN limit’, which is version specific:
– – The general soft limit is 16384 (16k) SCNs per second calculated by the number of second from Januari 1st, 1988 times 16384.
– – From version 12.2 on, with compatibility set to 12.2 or higher, the limit is 98304 (96k) SCNs per second calculated by the number of seconds from Januari 1st, 2008 times 98304.
– The RSL can not be exceeded, if a database session tries to go beyond the soft limit, an ORA 600 [2252] is raised and the transaction is rolled back. This will not corrupt data (but obviously the data in the transaction is not applied).
– The difference between the current SCN and the RSL SCN is known as ‘SCN headroom’.
– There have been bugs that can increase SCNs at a higher rate, decreasing the SCN headroom or even reaching the soft limit, but these have all been solved in the Januari 2012 CPU/PSU/patch bundles.
– When databases communicate which each other via a database link, the SCNs of both databases are synchronised by picking the highest of the two.

But it doesn’t really explain why patches must be applied before June 2019. However, another article is more concrete on that: Recommended patching and actions for Oracle database versions 12.1.0.1, 11.2.0.3 and earlier – before June 2019 (Doc ID 2335265.1). The first interesting thing mentioned is:

3. What is the change introduced by the patches listed above?
These patches increase the database’s current maximum SCN (system change number) limit.
At any point in time, the Oracle Database calculates a “not to exceed” limit for the number of SCNs a database can have used, based on the number of seconds elapsed since 1988.

So, this means the patched database have a change (increase) in the RSL.

These recommended patches enable the databases to allow for a higher current maximum SCN limit. The rate at which this limit is calculated can be referred to as the “SCN rate” and these patches help allow higher SCN rates to enable databases to support many times higher transaction rates than earlier releases.

And this means the RSL is increased from the 16k per second since Januari 1988 to the 96k per second since Januari 2008.

Please note that the patches only increase the max limit but the current SCN is not impacted. So, if all your databases don’t have any major change in transaction rate, the current SCN would still remain below the current maximum SCN limit and database links between newer (or patched) and unpatched databases would continue to work. The patches provide the safety measure to ensure that you don’t have any issue with dblinks independent of any possible future change in your transaction rate.

With the patches applied, this change in current maximum SCN limit will happen automatically starting 23rd June 2019.

That is important information! So with the patch applied to some databases and not applied to some other databases and patched and non-patched databases have database links between them, everything should remain working. This is true for any database at this moment, because the change of the limit will happen on the 23rd of June 2019, which at the time of writing is in the future.

Once the change of limit has happened at the 23rd of June 2019, database links between older, non-patched versions of the database and newer or patched versions of the database could be affected if after the 23rd of June 2019 the SCN rate is increased in a newer or patched database and it runs beyond the headroom available in databases with the 16k rate, because the earlier mentioned SCN synchronisation then will fail because it’s beyond the 16k rate database headroom meaning it can not synchronise with the newer dataebase.

So the problem we are talking about here is two databases talking with each other over a database link, which have a different RSL, which could lead to a situation that one database is at an SCN number which is too high for the other older or non-patched database, meaning the communication will fail, which will persist until the older or non-patched databases is able to reach the SCN number of the newer database over time, if that is possible at all.

A thing that is not clear to me at this point: the documentation to me seems to suggest that Oracle version 12.2 with compatibility set to 12.2 or higher versions already allows the higher number of SCNs per second. (the above point: – – From version 12.2 on, with compatibility set to 12.2 or higher, the limit is 98304 (96k) SCNs per second calculated by the number of seconds from Januari 1st, 2008) If that is true, the issue that is warned for could potentially already happen!

Luckily, there is are ways to investigate this:

The reported “newer” versions and the versions that are patched for the rate change have an SGA variable “kcmscnc_” that lists the current SCN rate of the database. There are 3 values that I have seen for “kcmscnc_”:
1: This is the rate of 16k per second since Januari 1st 1988.
2: ?
3: This is the rate of 96k per second since Januari 1st 2008.

SQL> select ksmfsnam, ksmfsadr, ksmfssiz from x$ksmfsv where ksmfsnam like 'kcmscnc_';
KSMFSNAM							 KSMFSADR	    KSMFSSIZ
---------------------------------------------------------------- ---------------- ----------
kcmscnc_							 000000006001579C	   4
SQL> oradebug peek 0x6001579C 4
[06001579C, 0600157A0) = 00000001

So, this databases is capable of switching (because kcmscnc_ exists), and is currently running with the 16k per second threshold.

In fact, I tested this on 11.2.0.4, 12.1.0.2 and 18.3.0.0, all these version report currently (currently is before June 2019) “1” or “scheme 1”. So the above mentioned rate of 96k for 12.2 and above does currently not apply for the soft limit or SCN headroom calculation for any database, including 12.2 and 18.3.

There is a less “hardcore” way to obtain this information, via the DBMS_SCN. This package equally to the “kcmscnc_” variable only exists if the database is of a version or patch version high enough to switch:

declare
  v_rsl number;
  v_headroom_in_scn number;
  v_headroom_in_sec number;
  v_cur_scn_compat number;
  v_max_scn_compat number;
begin
  dbms_scn.getcurrentscnparams(v_rsl, v_headroom_in_scn, v_headroom_in_sec, v_cur_scn_compat, v_max_scn_compat);
  dbms_output.put_line('reasonable scn limit (soft limit): '||to_char(v_rsl,'999,999,999,999,999,999'));
  dbms_output.put_line('headroom in scn                  : '||to_char(v_headroom_in_scn,'999,999,999,999,999,999'));
  dbms_output.put_line('headroom in sec                  : '||v_headroom_in_sec);
  dbms_output.put_line('current scn compatibility scheme : '||v_cur_scn_compat);
  dbms_output.put_line('max scn compatibility scheme     : '||v_max_scn_compat);
end;
/

SQL> /
reasonable scn limit (soft limit):	 16,104,861,483,008
headroom in scn 		 :	 16,104,860,802,471
headroom in sec 		 : 982962695
current scn compatibility scheme : 1
max scn compatibility scheme	 : 3

PL/SQL procedure successfully completed.

This is executed in a version 18.3.0.0 database. So the conclusion here is that currently all versions up to 18.3.0.0 are still compatible, because they all use the same SCN limit per second, which is referred to as ‘scheme 1’. However, on June 23, 2019 newer versions of the database will switch to a new scheme, which is referred to as ‘scheme 3’, which allows an more headroom.

First of all, I hope a lot of databases out there have enough headroom in the first place, and a modest rate of SCNs it is using per second, which means it doesn’t advance into the RSL. In such a case, when you got older versions that can not switch to scheme 3 and newer versions that will, I can see no reason to worry at all.

The second case is when your database is close to running out of headroom currently in scheme 1, and the SCN rate in the database is also close the limit, so you truly should worry when your database switches to scheme 3, it might actually run over the 16k per second limit, and when it does that long enough to run over RSL of scheme 1, communication over a database link between the current scheme 3 database with a scheme 1 database will be disrupted until the scheme 3 database SCN is low enough again to fit the scheme 1 RSL. If the SCN rate persists in the scheme 3 database, communication will be impossible between the scheme 1 and a scheme 3 database.

So, at this point it should be clear that the most important statistic for determining issues between scheme 1 and after June 2019 scheme 3 databases is the current headroom in your databases. For all of the databases involved that will have a database link between a scheme 1 and a scheme 3 database, you should investigate their SCN number and SCN rate. If some of these databases have run into the soft limit ora-600, ora 600 [2252], these are potential candidates for running over the soft limit when they switch to scheme 3.

However, the most important statistic is to see how close the current SCN is to the scheme 1 soft limit. This can be done using the following SQL (this SQL does not need a newer version of the database, and is tested with version 11.2.0.2 and higher):

col "RSL scheme 1" format 9,999,999,999,999,999
col "current value" format 9,999,999,999,999,999
select dbms_flashback.get_system_change_number "current value",
       ((((to_number(to_char(sysdate,'YYYY'))-1988)*12*31*24*60*60) +
       ((to_number(to_char(sysdate,'MM'))-1)*31*24*60*60) +
       (((to_number(to_char(sysdate,'DD'))-1))*24*60*60) +
       (to_number(to_char(sysdate,'HH24'))*60*60) +
       (to_number(to_char(sysdate,'MI'))*60) +
       (to_number(to_char(sysdate,'SS')))) * (16*1024)) "RSL scheme 1",
       round(dbms_flashback.get_system_change_number/((((to_number(to_char(sysdate,'YYYY'))-1988)*12*31*24*60*60) +
       ((to_number(to_char(sysdate,'MM'))-1)*31*24*60*60) +
       (((to_number(to_char(sysdate,'DD'))-1))*24*60*60) +
       (to_number(to_char(sysdate,'HH24'))*60*60) +
       (to_number(to_char(sysdate,'MI'))*60) +
       (to_number(to_char(sysdate,'SS')))) * (16*1024))*100,5) "% to RSL scheme 1"
from dual;

If a database is close the RSL (roughly speaking beyond 90%-95%), the next thing to do is measure if the database keeps on using SCNs and keeps on being close to the RSL. If that is true, an additional increase in SCN usage could in the current situation using scheme 1 lead to an ora-600 [2252], but if that database has switched to scheme 3 after June 2019, there will not be anything keeping that database from going beyond an SCN number that will exceed the RSL of scheme 1, which will then cause issues if that database has a database link with a scheme 1 database.

Is there anything you can do if you suspect or know a database will go over the scheme 1 RSL? Purely for this issue, the obvious solution would be to make sure you are on a version that will switch to scheme 3 on June 2019, so at least after June 2019 it will not run into ora-600 [2252].

However, if such a scheme 3 database needs to connect to an older scheme 1 database, you have two choices:
1. Potentially run over the the scheme 1 limit and disrupt the database link communication.
2. Stop a newer database from switching to scheme 3, potentially disrupt changes in the current database, but it guarantees the database link will always work.

To look into the switch to scheme 3, which oracle calls ‘auto rollover’, the following SQL can be used:

declare
  v_autorollover_date date;
  v_target_compat number;
  v_is_enabled boolean;
begin
  dbms_scn.getscnautorolloverparams(v_autorollover_date, v_target_compat, v_is_enabled);
  dbms_output.put_line('auto rollover date      : '||to_char(v_autorollover_date,'YYYY-MM-DD'));
  dbms_output.put_line('target scheme	        : '||v_target_compat);
  dbms_output.put_line('rollover enabled (1=yes): '||sys.diutil.bool_to_int(v_is_enabled));
end;
/
SQL> /
auto rollover date	: 2019-06-23
target scheme		: 3
rollover enabled (1=yes): 1

PL/SQL procedure successfully completed.

If you want to prevent a database from rolling over to scheme 3, the procedure dbms_scn.disableautorollover can be used:

exec dbms_scn.disableautorollover;

Obviously, the procedure enableautorollover does the opposite. Please mind to contact Oracle support before doing this with your production database, this is an undocumented procedure at this time.

Also mind that if you create a new database after June 23, 2019, with a new or patched version that can switch to scheme 3, it will probably be running scheme 3 by default. If you want to be absolutely sure it will not exceed the scheme 1 limit, you can revert it to scheme 1 manually using the alter database set scn compatibility N command in mount mode:

SQL> startup mount;
ORACLE instance started.

Total System Global Area 1048574496 bytes
Fixed Size		    8665632 bytes
Variable Size		  281018368 bytes
Database Buffers	  616562688 bytes
Redo Buffers		  142327808 bytes
Database mounted.
SQL> alter database set scn compatibility 1;

Database altered.

SQL> alter database open;

Database altered.

For this too I would strongly advise to contact Oracle support first. The purpose of this blogpost is to define the problem, show all the technical details that have to do with it, and show all the tools that are part of it. There is in no way anything in this article to tell you what you should do, it just shows everything that surrounds the switch to scheme 3 in June 2019.

Another view that might be beneficial is x$kcmscn. This view seems to be created to help looking if a scheme 3 database can connect to a scheme 1 database:

col cur_max_scn format 999,999,999,999,999
col pre_11_2_0_2_cur_max_scn format 999,999,999,999,999
select * from x$ksmscn;
ADDR		       INDX    INST_ID	   CON_ID CUR_MAX_RATE	  CUR_SCN
---------------- ---------- ---------- ---------- ------------ ----------
	 CUR_MAX_SCN DIFF_IN_DAYS DIFF_IN_HOURS PRE_11_2_0_2_CUR_MAX_SCN
-------------------- ------------ ------------- ------------------------
PRE_11_2_0_2_DIFF_IN_DAYS
-------------------------
00007F773DEDAE10	  0	     1		0	     0	   800918
  16,108,830,064,640	    11379	 273112       16,108,830,064,640
		186444783

Because currently (before June 2019) every database by default will be in scheme 1, the cur_max_scn and pre_11_2_0_2_cur_max_scn are identical. I even believe the column naming is wrong, the first version that can switch if it is patched to a high enough PSU version is 11.2.0.3, I do believe the column name is suggesting scheme 1 databases are databases of a version lower than 11.2.0.2, not including 11.2.0.2.

Conclusion.
I think there’s been a lot of fuzz for something that in most cases is not an issue. This article is supposed to give you all the knowledge and the tools to determine how it looks like in your situation.

This might be an issue if you happen to have one or more databases that are high on SCN numbering, and continues to take a lot of SCN numbers, and will be converted to a scheme 3 database on June 29, 2019 and is suspected to increase on taking SCN numbers for whatever reason AND it has a database link to a scheme 1 database that remains scheme 1. That’s a lot of ifs.

On the other hand you only need one database to be high in SCN numbering which continues to take a lot of SCNs keeping it close to the soft limit, which will propagate its SCN to other databases if it is linked, or the required properties of the problem spread out over multiple linked databases.

Again, I do not advise anything in this article, the purpose here is to provide all the details that surround it so you can make the best decision for yourself.

A look into oracle redo, part 11: log writer worker processes

April 30, 2018

Exadata, Linux, Networking, Oracle, Oracle EE, Oracle XE

5 Comments

Starting from Oracle 12, in a default configured database, there are more log writer processes than the well known ‘LGWR’ process itself, which are the ‘LGnn’ processes:

$ ps -ef | grep test | grep lg
oracle   18048     1  0 12:50 ?        00:00:13 ora_lgwr_test
oracle   18052     1  0 12:50 ?        00:00:06 ora_lg00_test
oracle   18056     1  0 12:50 ?        00:00:00 ora_lg01_test

These are the log writer worker processes, for which the minimal amount is equal to the amount public redo strands. Worker processes are assigned to a group, and the group is assigned to a public redo strand. The amount of worker processes in the group is dependent on the undocumented parameter “_max_log_write_parallelism”, which is one by default.

The actual usage of the worker processes is dependent in the first place on the value of the undocumented parameter “_use_single_log_writer”, for which the default value is ‘ADAPTIVE’, which means it’s switching automatically between ‘single log writer mode’, which is the traditional way of the LGWR process handling everything that the log writer functionality needs to do, and the ‘scalable log writer mode’, which means the log writer functionality is presumably using the log writer worker processes.

Other values for “_use_single_log_writer” are ‘TRUE’ to set ‘single log writer mode’, or ‘FALSE’ to set ‘scalable log writer mode’ fixed.

I assume most readers of this blog will know that the master log writer idle work cycle is sleeping on a semaphore (semtimedop()) under the wait event ‘rdbms ipc message’ for 3 seconds, then performs some “housekeeping”, after which it’ll sleep again repeating the small cycle of sleeping and housekeeping. For the log writer worker processes, this looks different if you look at the wait event information of the log writer worker processes:

135,59779,@1    14346                    DEDICATED oracle@memory-presentation.local (LGWR)	    time:1909.44ms,event:rdbms ipc message,seq#:292
48,34282,@1     14350                    DEDICATED oracle@memory-presentation.local (LG00)	    time:57561.85ms,event:LGWR worker group idle,seq#:150
136,24935,@1    14354                    DEDICATED oracle@memory-presentation.local (LG01)	    time:112785.66ms,event:LGWR worker group idle,seq#:74

The master log writer process (LGWR) has been sleeping for 1.9s when I queried the database, and it will sleep for 3 seconds, and then do some work and sleep again. However, the log writer worker processes have been sleeping for much longer: LG00 for 57.6s and LG01 for 112.8s, and the event is different: ‘LGWR worker group idle’. How is this implemented? Let’s look!

$ strace -p $(pgrep lg01)
strace: Process 14354 attached
semtimedop(360448, [{27, -1, 0}], 1, {3, 0}) = -1 EAGAIN (Resource temporarily unavailable)
semtimedop(360448, [{27, -1, 0}], 1, {3, 0}) = -1 EAGAIN (Resource temporarily unavailable)

I used strace on the LG01 process, and it’s still doing the same as most idle background processes are doing: sleeping on a semaphore for 3 seconds. But, it does not end its wait like LGWR does, the event the log writer worker processes are waiting in keeps on being timed.

Using a pin tools debugtrace shows the following:

 | | < semtimedop+0x000000000023 returns: 0xffffffffffffffff
 | | > __errno_location(0x38000, 0x7ffce278c328, ...)
 | | | > fthread_self(0x38000, 0x7ffce278c328, ...)
 | | | < fthread_self+0x000000000024 returns: 0
 | | < __errno_location+0x000000000010 returns: 0x7f7e930a26a0
 | < sskgpwwait+0x00000000014e returns: 0
 < skgpwwait+0x0000000000e0 returns: 0
 > ksuSdiInProgress(0x19e80, 0x19e80, ...)
 < ksuSdiInProgress+0x000000000035 returns: 0
 > sltrgftime64(0x19e80, 0x19e80, ...)
 | > clock_gettime@plt(0x1, 0x7ffce278c3a0, ...)
 | | > clock_gettime(0x1, 0x7ffce278c3a0, ...)
 | | < clock_gettime+0x000000000069 returns: 0
 | < clock_gettime+0x00000000003a returns: 0
 < sltrgftime64+0x00000000004c returns: 0x19c253f3ff
 > kslwo_getcbk(0xa2, 0xd80fa62, ...)
 < kslwo_getcbk+0x000000000017 returns: 0
 > kgslwait_last_waitctx_time_waited_usecs(0x7f7e930a29a0, 0x6dfd01c0, ...)
 < kgslwait_last_waitctx_time_waited_usecs+0x000000000045 returns: 0x25e5e80
 > kskiorm(0x6d1854a8, 0, ...)
 < kskiorm+0x00000000001e returns: 0
 > kfias_iswtgon_ksfd(0x6d1854a8, 0, ...)
 < kfias_iswtgon_ksfd+0x00000000002b returns: 0
 > kxdbio_has_work(0x7ffce278c3c4, 0x6003d010, ...)
 < kxdbio_has_work+0x000000000027 returns: 0
 > skgpwwait(0x7ffce278c630, 0x7f7e930a7ca0, ...)
 | > kslwait_conv_wait_time(0x2dc6c0, 0x7f7e930a7ca0, ...)
 | < kslwait_conv_wait_time+0x000000000027 returns: 0x2dc6c0
 | > sskgpwwait(0x7ffce278c630, 0x7f7e930a7ca0, ...)
 | | > semtimedop(0x38000, 0x7ffce278c328, ...)
 | | < semtimedop+0x000000000023 returns: 0xffffffffffffffff

And a full stack trace of a log writer worker look like this:

$ pstack $(pgrep lg01)
#0  0x00007feda8eaebda in semtimedop () at ../sysdeps/unix/syscall-template.S:81
#1  0x0000000010f9cca6 in sskgpwwait ()
#2  0x0000000010f9a2e8 in skgpwwait ()
#3  0x0000000010a66995 in ksliwat ()
#4  0x0000000010a65d25 in kslwaitctx ()
#5  0x00000000031fb4d0 in kcrfw_slave_queue_remove ()
#6  0x00000000031fad2a in kcrfw_slave_group_main ()
#7  0x00000000012160fa in ksvrdp_int ()
#8  0x000000000370d99a in opirip ()
#9  0x0000000001eb034a in opidrv ()
#10 0x0000000002afedf1 in sou2o ()
#11 0x0000000000d0547a in opimai_real ()
#12 0x0000000002b09b31 in ssthrdmain ()
#13 0x0000000000d05386 in main ()

If you combine the pstack backtrace and the debugtrace information, you see that the idle cycle does not leave the ‘ksliwat’ function, so the wait event is not finished. Quickly looking at the other functions, it’s easy to spot it reads the system clock (sltrgftime64), updates some information (kgslwait_last_waitctx_time_waited_usecs) and then performs some proactive IO checks (kskiorm, kfias_iswtgon_ksfd, kxdbio_has_work) after which it calls the post/wait based functions to setup the semaphore again.

Conclusion so far is the log writer workers do perform a 3 second sleep just like the master log writer, however the wait event ‘LGWR worker group idle’ is not interrupted like ‘rdbms ipc message’ is for the master log writer. This means the wait time for the event for each worker process indicates the last time the worker process actually performed something. A next logical question then is: but what do the log writer worker processes perform? Do they entirely take over the master log writer functionality, or do they work together with the master log writer?

In order to fully understand the next part, it is very beneficial to read up on how the log writer works in ‘single log writer’ mode, where the master log writer handling the idle and work cycle itself:
– https://fritshoogland.wordpress.com/2018/02/20/a-look-into-into-oracle-redo-part-4-the-log-writer-null-write/
– https://fritshoogland.wordpress.com/2018/02/27/a-look-into-oracle-redo-part-5-the-log-writer-writing/

If you want to perform this investigation yourself, make sure the database is in ‘scalable log writer’ mode, by setting “_use_single_log_writer” to FALSE. This is exactly what I did in order to make sure a log write is done in ‘scalable log writer’ mode.

Now let’s first apply some logic. Above the idle cycle of a log writer worker process is shown. Based on the ‘log writer null write’ blog post, we know that the log writer does advance the LWN and On-disk SCN every 3 seconds. Clearly, the log writer worker process does not do that. So that must mean the master log writer is still performing that function. It would also make very much sense, because it doesn’t matter for scalability if the master log writer performs the function of advancing the LWN and On-disk SCN or a worker process, nothing is waiting on it. Plus, if the master log writer performs most of its functions just like in ‘single log writer’ mode, the change to scalable mode would mean no change for client processes, any committing process must semop() the log writer to start writing.

Let’s look at the relevant debugtrace output of the master log writer in scalable log writer mode:

 | > kcrfw_redo_write_driver(0, 0, ...)
 | | > kcrfw_handle_member_write_errors(0, 0, ...)
 | | < kcrfw_handle_member_write_errors+0x000000000020 returns: 0x600161a0
 | | > kcmgtsf(0, 0, ...)
 | | | > sltrgatime64(0, 0, ...)
 | | | | > sltrgftime64(0, 0, ...)
 | | | | | > clock_gettime@plt(0x1, 0x7fff1fe13010, ...)
 | | | | | | > clock_gettime(0x1, 0x7fff1fe13010, ...)
 | | | | | | < clock_gettime+0x000000000069 returns: 0
 | | | | | < clock_gettime+0x00000000003a returns: 0
 | | | | < sltrgftime64+0x00000000004c returns: 0x53747fe42
 | | | < sltrgatime64+0x00000000003e returns: 0x155d4fd
 | | < kcmgtsf+0x00000000032f returns: 0x3a182314
 | | > kcrfw_slave_adaptive_updatemode(0, 0x600161a0, ...)
 | | < kcrfw_slave_adaptive_updatemode+0x000000000080 returns: 0x7efe34d1f760
 | | > kcrfw_defer_write(0, 0x600161a0, ...)
 | | < kcrfw_defer_write+0x000000000038 returns: 0x7efe34d1f760
 | | > kcrfw_slave_queue_find(0, 0x600161a0, ...)
 | | < kcrfw_slave_queue_find+0x0000000000f1 returns: 0
 | | > kcrfw_slave_queue_setpreparing(0, 0x1, ...)
 | | < kcrfw_slave_queue_setpreparing+0x000000000021 returns: 0
 | | > kcrfw_slave_group_switchpic(0, 0x1, ...)
 | | < kcrfw_slave_group_switchpic+0x000000000050 returns: 0x699b4508
 | | > skgstmGetEpochTs(0, 0x1, ...)
 | | | > gettimeofday@plt(0x7fff1fe13070, 0, ...)
 | | | < __vdso_gettimeofday+0x0000000000fe returns: 0
 | | < skgstmGetEpochTs+0x000000000049 returns: 0x20debfd6192e5
 | | > kcsnew3(0x600113b8, 0x7fff1fe13228, ...)
 | | | > kcsnew8(0x600113b8, 0x7fff1fe13070, ...)
 | | | | > kslgetl(0x60049800, 0x1, ...)
 | | | | < kslgetl+0x00000000012f returns: 0x1
 | | | | > kslfre(0x60049800, 0x1, ...)
 | | | | < kslfre+0x0000000001e2 returns: 0
 | | | < kcsnew8+0x000000000117 returns: 0
 | | | > ub8_to_kscn_impl(0x66c3c7, 0x7fff1fe13228, ...)
 | | | < ub8_to_kscn_impl+0x000000000031 returns: 0
 | | < kcsnew3+0x00000000006f returns: 0x8000
 | | > ktfwtsm(0x3a182314, 0x7fff1fe13228, ...)
 | | | > kcmgtsf(0x2, 0x7fff1fe13228, ...)
 | | | | > sltrgatime64(0x2, 0x7fff1fe13228, ...)
 | | | | | > sltrgftime64(0x2, 0x7fff1fe13228, ...)
 | | | | | | > clock_gettime@plt(0x1, 0x7fff1fe12fe0, ...)
 | | | | | | | > clock_gettime(0x1, 0x7fff1fe12fe0, ...)
 | | | | | | | < clock_gettime+0x000000000069 returns: 0
 | | | | | | < clock_gettime+0x00000000003a returns: 0
 | | | | | < sltrgftime64+0x00000000004c returns: 0x537484a6d
 | | | | < sltrgatime64+0x00000000003e returns: 0x155d511
 | | | < kcmgtsf+0x0000000001b2 returns: 0x3a182314
 | | | > kcmtdif(0x3a182314, 0x3a182314, ...)
 | | | < kcmtdif+0x00000000001b returns: 0
 | | | > ksl_get_shared_latch_int(0x60050340, 0x6ddb1408, ...)
 | | | < ksl_get_shared_latch_int+0x00000000016b returns: 0x1
 | | <> kslfre(0x60050340, 0x66c3c7, ...)
 | | < kslfre+0x0000000001e2 returns: 0
 | | > kcn_stm_write(0x7fff1fe13228, 0x66c3c7, ...)
 | | | > kstmgetsectick(0x7fff1fe13228, 0x66c3c7, ...)
 | | | < kstmgetsectick+0x00000000003a returns: 0x5ae4c494
 | | | > ksl_get_shared_latch_int(0x6004ee40, 0x6ddb1408, ...)
 | | | < ksl_get_shared_latch_int+0x00000000016b returns: 0x1
 | | <> kslfre(0x6004ee40, 0x2244, ...)
 | | < kslfre+0x0000000001e2 returns: 0
 | | > kcrfw_redo_write_initpic(0x699b4508, 0x7fff1fe13228, ...)
 | | | > kscn_to_ub8_impl(0x7fff1fe13228, 0x7fff1fe13228, ...)
 | | | < kscn_to_ub8_impl+0x00000000003e returns: 0x66c3c7
 | | < kcrfw_redo_write_initpic+0x0000000000dc returns: 0x3a182314
 | | > kscn_to_ub8_impl(0x7fff1fe13228, 0, ...)
 | | < kscn_to_ub8_impl+0x00000000003e returns: 0x66c3c7
 | | > kcrfw_gather_lwn(0x7fff1fe13268, 0x699b4508, ...)
 | | | > kslgetl(0x6abe4538, 0x1, ...)
 | | | < kslgetl+0x00000000012f returns: 0x1
 | | | > kcrfw_gather_strand(0x7fff1fe13268, 0, ...)
 | | | < kcrfw_gather_strand+0x0000000000c2 returns: 0
 | | | > kslfre(0x6abe4538, 0x17d5f, ...)
 | | | < kslfre+0x0000000001e2 returns: 0
 | | | > kslgetl(0x6abe45d8, 0x1, ...)
 | | | < kslgetl+0x00000000012f returns: 0x1
 | | | > kcrfw_gather_strand(0x7fff1fe13268, 0x1, ...)
 | | | < kcrfw_gather_strand+0x0000000000c2 returns: 0
 | | | > kslfre(0x6abe45d8, 0x137, ...)
 | | | < kslfre+0x0000000001e2 returns: 0
 | | < kcrfw_gather_lwn+0x00000000065c returns: 0xffffffff
 | | > krsh_trace(0x1000, 0x200, ...)
 | | < krsh_trace+0x00000000005d returns: 0
 | | > kspgip(0x71e, 0x1, ...)
 | | < kspgip+0x00000000023f returns: 0
 | | > kcrfw_slave_queue_setpreparing(0, 0, ...)
 | | < kcrfw_slave_queue_setpreparing+0x000000000021 returns: 0
 | | > kcrfw_slave_queue_flush_internal(0x1, 0, ...)
 | | < kcrfw_slave_queue_flush_internal+0x0000000000d7 returns: 0x1
 | | > kcrfw_do_null_write(0, 0, ...)
 | | | > kcrfw_slave_phase_batchdo(0, 0, ...)
 | | | | > kcrfw_slave_phase_enter(0, 0x9b, ...)
 | | | | < kcrfw_slave_phase_enter+0x000000000449 returns: 0
 | | | <> kcrfw_slave_phase_exit(0, 0x9b, ...)
 | | | < kcrfw_slave_phase_exit+0x00000000035a returns: 0
 | | | > kcrfw_post(0, 0, ...)
 | | | | > kcrfw_slave_single_getactivegroup(0, 0, ...)
 | | | | < kcrfw_slave_single_getactivegroup+0x000000000047 returns: 0x6a9a0718
 | | | | > kspGetInstType(0x1, 0x1, ...)
 | | | | | > vsnffe_internal(0x19, 0x1, ...)
 | | | | | | > vsnfprd(0x19, 0x1, ...)
 | | | | | | < vsnfprd+0x00000000000f returns: 0x8
 | | | | | | > kfIsASMOn(0x19, 0x1, ...)
 | | | | | | <> kfOsmInstanceSafe(0x19, 0x1, ...)
 | | | | | | < kfOsmInstanceSafe+0x000000000031 returns: 0
 | | | | | < vsnffe_internal+0x0000000000a7 returns: 0
 | | | | | > kspges(0x115, 0x1, ...)
 | | | | | < kspges+0x00000000010f returns: 0
 | | | | < kspGetInstType+0x0000000000b1 returns: 0x1
 | | | | > kcrfw_slave_phase_enter(0x1, 0x9b, ...)
 | | | | < kcrfw_slave_phase_enter+0x00000000006f returns: 0x9b
 | | | | > kcscu8(0x60016290, 0x7fff1fe12f98, ...)
 | | | | < kcscu8+0x000000000047 returns: 0x1
 | | | | > kcsaj8(0x60016290, 0x7fff1fe12f38, ...)
 | | | | < kcsaj8+0x0000000000dc returns: 0x1
 | | | | > kcrfw_slave_phase_exit(0x1, 0x9b, ...)
 | | | | < kcrfw_slave_phase_exit+0x00000000008e returns: 0
 | | | | > kslpsemf(0x97, 0, ...)
 | | | | | > ksl_postm_init(0x7fff1fe0ac30, 0x7fff1fe12c50, ...)
 | | | | | < ksl_postm_init+0x00000000002b returns: 0
 | | | | < kslpsemf+0x0000000006b5 returns: 0x1f
 | | | | > kcrfw_slave_barrier_nonmasterwait(0x6a9a0720, 0x4, ...)
 | | | | < kcrfw_slave_barrier_nonmasterwait+0x000000000035 returns: 0x600161a0
 | | | < kcrfw_post+0x000000000c1c returns: 0xd3
 | | < kcrfw_do_null_write+0x0000000000b2 returns: 0xd3
 | < kcrfw_redo_write_driver+0x000000000535 returns: 0xd3

The highlighted functions are extra functions executed when the instance is set to scalable log writer mode, or when adaptive mode has set the instance to scalable log writer mode. This means that the changes between the modes is minimal when there’s no writes, and outside of a few extra functions, the log writer does exactly the same.

The absence of any spectacular changes in the behaviour of the log writer when in scalable log writer mode when there are no writes does hint what the actual changes will be of the scalable mode, which is how writing is handled. In single log writer mode, the most time the log writer is process is likely to spend on is writing the change vectors into the online redologfiles, and maybe, if you have a bad application (!) semop()-ing foreground sessions will be second, if there are a large number of processes committing, because every process needs to be semop()-ed individually. These two functions, along with some other functionality are exactly what the log writer worker processes are doing.

This means that foreground processes do nothing different in scalable log writer mode, they signal (semop) the master log writer, which will investigate the public redo strands, and if the master log writer finds change vectors to write, it will assign log writer worker processes to perform the write, and the log writer worker process will semop() the foreground sessions to indicate the redo has been written when the instance is in post/wait mode, or do not semop() when the instance is in polling mode.

This is the entire function flow of a write when the instance is in scalable log writer mode:

 | > kcrfw_slave_queue_insert(0, 0xd3, ...)
 | | > kcrfw_slave_group_setcurrsize(0, 0, ...)
 | | < kcrfw_slave_group_setcurrsize+0x0000000001d1 returns: 0x1
 | | > _intel_fast_memcpy(0x6a9a05f8, 0x7ffdae335fa0, ...)
 | | <> _intel_fast_memcpy.P(0x6a9a05f8, 0x7ffdae335fa0, ...)
 | | <> __intel_ssse3_rep_memcpy(0x6a9a05f8, 0x7ffdae335fa0, ...)
 | | < __intel_ssse3_rep_memcpy+0x000000002798 returns: 0x6a9a05f8
 | | > kcrfw_slave_group_postall(0, 0xf0, ...)
 | | | > ksvgcls(0, 0xf0, ...)
 | | | < ksvgcls+0x000000000021 returns: 0
 | | | > ksl_post_proc(0x6ddb32f0, 0, ...)
 | | | <> kskpthr(0x6ddb32f0, 0, ...)
 | | | <> kslpsprns(0x6ddb32f0, 0, ...)
 | | | | > ksl_update_post_stats(0x6ddb32f0, 0, ...)
 | | | | | > dbgtTrcData_int(0x7f464c0676c0, 0x2050031, ...)
 | | | | | | > dbgtBucketRedirect(0x7f464c0676c0, 0x7ffdae335338, ...)
 | | | | | | < dbgtBucketRedirect+0x000000000050 returns: 0x1
 | | | | | | > dbgtIncInMemTrcRedirect(0x7f464c0676c0, 0x6fa, ...)
 | | | | | | < dbgtIncInMemTrcRedirect+0x000000000035 returns: 0x1
 | | | | | | > skgstmGetEpochTs(0x7f464c0676c0, 0x6fa, ...)
 | | | | | | | > gettimeofday@plt(0x7ffdae334e40, 0, ...)
 | | | | | | | < __vdso_gettimeofday+0x0000000000fe returns: 0
 | | | | | | < skgstmGetEpochTs+0x000000000049 returns: 0x20e067375b55d
 | | | | | | > dbgtrRecAllocate(0x7f464c0676c0, 0x7ffdae3352e0, ...)
 | | | | | | | > dbgtrPrepareWrite(0x7f464c0676c0, 0x65accba0, ...)
 | | | | | | | < dbgtrPrepareWrite+0x00000000011c returns: 0x4
 | | | | | | < dbgtrRecAllocate+0x000000000144 returns: 0x1
 | | | | | | > _intel_fast_memcpy(0x65acda30, 0x7ffdae3353d8, ...)
 | | | | | | <> _intel_fast_memcpy.P(0x65acda30, 0x7ffdae3353d8, ...)
 | | | | | | <> __intel_ssse3_rep_memcpy(0x65acda30, 0x7ffdae3353d8, ...)
 | | | | | | < __intel_ssse3_rep_memcpy+0x000000002030 returns: 0x65acda30
 | | | | | | > dbgtrRecEndSegment(0x7f464c0676c0, 0x7ffdae3352e0, ...)
 | | | | | | < dbgtrRecEndSegment+0x00000000011c returns: 0x77c000a4
 | | | | | < dbgtTrcData_int+0x000000000323 returns: 0x77c000a4
 | | | | < ksl_update_post_stats+0x00000000024f returns: 0x77c000a4
 | | | | > skgpwpost(0x7ffdae335480, 0x7f464c0acca0, ...)
 | | | | <> sskgpwpost(0x7ffdae335480, 0x7f464c0acca0, ...)
 | | | | | > semop@plt(0xc0000, 0x7ffdae335410, ...)
 | | | | | < semop+0x00000000000f returns: 0
 | | | | < sskgpwpost+0x00000000009a returns: 0x1
 | | | < kslpsprns+0x0000000001c3 returns: 0
 | | < kcrfw_slave_group_postall+0x0000000000a8 returns: 0
 | < kcrfw_slave_queue_insert+0x0000000001b6 returns: 0x667bc540

After the instance has established there are change vectors in kcrfw_gather_lwn, in single log writer mode, the function kcrfw_redo_write is called, which will call kcrfw_do_write which handles the writing, and kslpslf to semop any waiting processes among other things. Now in scalable log writer mode, kcrfw_slave_queue_insert is called which assigns work to worker processes, and then kcrfw_slave_group_postall is called to semop one or more worker processes.

The worker processes are sleeping on a semaphore, and if a process gets signalled, it exits the kcrfw_slave_queue_remove function, ends the wait event, and calls kcrfw_redo_write, just like the master log writer process would call in single log writer mode, which includes doing the write (kcrfw_do_write) and posting the foregrounds (kslpslf), exactly all the functions.

Conclusion.
The adaptive scalable log writer processes function has been silently introduced with Oracle 12, although a lot of the used functionality has been available more or less in earlier versions. It is a fully automatic feature which will turn itself on and off based on heuristics. The purpose of this article is to explain how it works and what it is doing. Essentially, all the functionality that surrounds a log writer write has been moved to a worker process, which means the work can be done in parallel with multiple processes, whilst all the work outside of the work around the write, which is not performance critical, is left with the master log writer.

A look into oracle redo: index and overview

April 17, 2018

Exadata, Linux, Oracle, Oracle EE, Oracle XE

A look into oracle redo, part 10: commit_wait and commit_logging

April 16, 2018

Exadata, Oracle, Oracle EE, Oracle XE

1 Comment

The redo series would not be complete without writing about changing the behaviour of commit. There are two ways to change commit behaviour:

1. Changing waiting for the logwriter to get notified that the generated redo is persisted. The default is ‘wait’. This can be set to ‘nowait’.
2. Changing the way the logwriter handles generated redo. The default is ‘immediate’. This can be set to ‘batch’.

There are actually three ways these changes can be made:
1. As argument of the commit statement: ‘commit’ can be written as ‘commit write wait immediate’ (statement level).
2. As a system level setting. By omitting an explicit commit mode when executing the commit command, the setting as set with the parameters commit_wait (default: wait) and commit_logging (default: immediate).
3. As a session level setting. By omitting an explicit commit mode, but by setting either commit_wait or commit_logging it overrides the settings at the system level.

At this point I should say that in my personal opinion, if you need to change this, there is something very wrong with how the database is used in the first place. This can enhance performance a bit (totally depending on what you are doing and how your hardware looks like), but it does nothing magic, as you will see.

a) commit wait/nowait
I ran a pin tools debugtrace on a session that commits explicitly with the write mode explicitly set to wait (the default), and a session that commits explicitly with the write mode set to nowait. If you took the time to read the other redo related articles you know that a commit generates changes vectors that are written in the public redo strand, changes the transaction table in the undo segment header and then signals the logwriter to write in kcrf_commit_force_int, releases all transactional control on the rows in the transaction that are committed, after which kcrf_commit_force_int is called again in order to wait for the logwriter to get notified that the change vectors have been persisted.

When commit is set to nowait, actually what happens is very simple: everything that is executed in ‘wait mode’ commit is executed in ‘nowait mode’ too, except for calling the kcrf_commit_force_int a second time, which is the functionality to wait for the notification from the logwriter.

commit wait:

 | | < kpoal8+0x000000000f8c returns: 0x2
 | | > ksupop(0x1, 0x7a87a9a0, ...)
 | | | > ksugit_i(0x11526940, 0x7a87a9a0, ...)
 | | | < ksugit_i+0x00000000002a returns: 0
 | | | > _setjmp@plt(0x7ffda5959c50, 0x7a87a9a0, ...)
 | | | <> __sigsetjmp(0x7ffda5959c50, 0, ...)
 | | | <> __sigjmp_save(0x7ffda5959c50, 0, ...)
 | | | < __sigjmp_save+0x000000000025 returns: 0
 | | | > kcbdsy(0x7ffda5959c50, 0x7f3011cbc028, ...)
 | | | <> kcrf_commit_force_int(0x7f3011d75e10, 0x1, ...)
...
 | | | < kcrf_commit_force_int+0x000000000b9c returns: 0x1
 | | | > kslws_check_waitstack(0x3, 0x7f3011d82f40, ...)
 | | | < kslws_check_waitstack+0x000000000065 returns: 0
 | | | > kssdel(0x7a87a9a0, 0x1, ...)
 | | | | > kpdbUidToId(0, 0x1, ...)
 | | | | < kpdbUidToId+0x00000000014e returns: 0
 | | | | > kss_del_cb(0x7ffda5959b50, 0x7f3011d82f40, ...)
 | | | | | > kpdbUidToId(0, 0x7f3011d82f40, ...)
 | | | | | < kpdbUidToId+0x00000000014e returns: 0
 | | | | | > ksudlc(0x7a87a9a0, 0x1, ...)

commit nowait:

 | | < kpoal8+0x000000000f8c returns: 0x2
 | | > ksupop(0x1, 0x63c82a38, ...)
 | | | > ksugit_i(0x11526940, 0x63c82a38, ...)
 | | | < ksugit_i+0x00000000002a returns: 0
 | | | > _setjmp@plt(0x7fff43332a50, 0x63c82a38, ...)
 | | | <> __sigsetjmp(0x7fff43332a50, 0, ...)
 | | | <> __sigjmp_save(0x7fff43332a50, 0, ...)
 | | | < __sigjmp_save+0x000000000025 returns: 0
 | | | > kslws_check_waitstack(0x3, 0x7fd1cea22028, ...)
 | | | < kslws_check_waitstack+0x000000000065 returns: 0
 | | | > kssdel(0x63c82a38, 0x1, ...)
 | | | | > kpdbUidToId(0, 0x1, ...)
 | | | | < kpdbUidToId+0x00000000014e returns: 0
 | | | | > kss_del_cb(0x7fff43332950, 0x7fd1ceae8f40, ...)
 | | | | | > kpdbUidToId(0, 0x7fd1ceae8f40, ...)
 | | | | | < kpdbUidToId+0x00000000014e returns: 0
 | | | | | > ksudlc(0x63c82a38, 0x1, ...)

Yes, it’s that simple. In normal commit mode, commit wait, in ksupop (kernel service user pop (restore) user or recursive call) a call to kcbdsy is executed, which performs a tailcall to kcrf_commit_force_int. In nowait commit mode, kcbdsy is simply not called in ksupop, which actually exactly does what nowait means, the waiting for the logwriter notification is not done.

b) commit immediate/batch
I ran a pin tools debugtrace on a session that commits explicitly with the write mode explicitly set to immediate, and a session that commits explicitly with the write mode set to batch. If you read the other redo related articles you know that a commit generates changes vectors that are written in the public redo strand, changes the transaction table in the undo segment header and then signals the logwriter to write in kcrf_commit_force_int, then releases all transactional control on the rows in the transaction that are committed, after which kcrf_commit_force_int is called again in order to wait for the logwriter to get notified that the change vectors have been persisted.

When commit is set to batch, actually what happens is very simple: everything is done exactly the same in ‘immediate mode’ commit, except for calling the kcrf_commit_force_int the first time, which is the functionality that triggers the logwriter to write. So it looks like ‘batch mode’ is not explicitly batching writes for the logwriter, but rather the disablement of the signal to the logwriter to write right after the change vectors have been copied and the blocks are changed. But that is not all…

I noticed something weird when analysing the calls in the debugtrace of ‘commit write batch’: not only was the first invocation of kcrf_commit_force_int gone, the second invocation of kcrf_commit_force_int was also gone too! That is weird, because the Oracle documentation says:

WAIT | NOWAIT

Use these clauses to specify when control returns to the user.

The WAIT parameter ensures that the commit will return only after the corresponding redo is persistent in the online redo log. Whether in BATCH or IMMEDIATE mode, when the client receives a successful return from this COMMIT statement, the transaction has been committed to durable media. A crash occurring after a successful write to the log can prevent the success message from returning to the client. In this case the client cannot tell whether or not the transaction committed.

The NOWAIT parameter causes the commit to return to the client whether or not the write to the redo log has completed. This behavior can increase transaction throughput. With the WAIT parameter, if the commit message is received, then you can be sure that no data has been lost.

If you omit this clause, then the transaction commits with the WAIT behavior.

The important, and WRONG thing, is in the last line: ‘if you omit this clause, then the transaction commits with the WAIT behavior’. Actually, if the commit mode is set to batch, the commit wait mode flips to nowait with it. It does perform the ultimate batching, which is not sending a signal to the logwriter at all, so what happens is that change vectors in the public redo strands are written to disk by the logwriter only every 3 seconds, because that is the timeout for the logwriter sleeping on a semaphore, after which it obtains any potential redo to write via information in kcrfsg_ and KCRFA structures. This is important, because with NOWAIT behaviour, there is no guarantee changes have been persisted for the committing session.

I was surprised to find this, which for me it meant I was searching for ‘kcrf_commit_force_int’ in the debugtrace of a commit with the ‘write batch’ arguments, and did not find any of them. Actually, this has been reported by Marcin Przepiorowski in a comment on an article by Christian Antognini on this topic.

Can this commit batching be changed to include waiting for the logwriter? Yes, actually it can if you explicitly include ‘wait’ with the commit write batch. It is very interesting the kcrf_commit_force_int function then comes back at a totally different place:

 | | | | | | | | | | | | | < ktuulc+0x000000000119 returns: 0
 | | | | | | | | | | | | | > ktudnx(0x69fc8eb0, 0, ...)
 | | | | | | | | | | | | | | > ktuIMTabCacheCommittedTxn(0x69fc8eb0, 0x7ffe9eb79e74, ...)
 | | | | | | | | | | | | | | < ktuIMTabCacheCommittedTxn+0x000000000071 returns: 0
 | | | | | | | | | | | | | | > kslgetl(0x6ab9d6e8, 0x1, ...)
 | | | | | | | | | | | | | | < kslgetl+0x00000000012f returns: 0x1
 | | | | | | | | | | | | | | > kslfre(0x6ab9d6e8, 0x6ab9ce00, ...)
 | | | | | | | | | | | | | | < kslfre+0x0000000001e2 returns: 0
 | | | | | | | | | | | | | < ktudnx+0x0000000005e4 returns: 0
 | | | | | | | | | | | | | > ktuTempdnx(0x69fc8eb0, 0, ...)
 | | | | | | | | | | | | | < ktuTempdnx+0x000000000083 returns: 0
 | | | | | | | | | | | | | > kcb_sync_last_change(0x69fc8eb0, 0x6df64df8, ...)
 | | | | | | | | | | | | | <> kcrf_commit_force_int(0x7f525ba19c00, 0x1, ...)
...
 | | | | | | | | | | | | | < kcrf_commit_force_int+0x000000000b9c returns: 0x1
 | | | | | | | | | | | | | > kghstack_free(0x7f525bb359a0, 0x7f525690ead8, ...)
 | | | | | | | | | | | | | < kghstack_free+0x00000000005a returns: 0
 | | | | | | | | | | | | < ktucmt+0x000000000e0c returns: 0

Instead of simply keeping the separate call after the transaction in the ksupop function, described above with commit wait/nowait, which is kcrf_commit_force_int with second argument set to 1, which means it notifies the logwriter as well as waits for the logwriter notification of the write, it is now is called after the function to clear the TX enqueue (ktuulc) and the undo transaction count has been lowered (ktudnx) at the end of the ktucmt function as a tailcall of kcb_sync_last_change, which wasn’t called before. Of course this limits the IO batching opportunities.

Conclusion
Do not change your database or even your session to make your commit faster. If you must, read this article carefully and understand the trade offs. One trade off which hasn’t been highlighted is: this might change in a different version, and it requires some effort to investigate. And again: if you still are considering this: probably you have a different problem that you should look at. Do not take this option in desperation to hope for a magical restoration of performance.

The commit_write option nowait does trigger the logwriter to write (the first invocation of the kcrf_commit_force_int function), but it does not wait for write confirmation. The commit_logging option batch does something different than the documentation says it does, it does not issue a signal to the logwriter, nor wait for it. This way the logwriter can wait the full three seconds before it times out on its semaphore and write what is in the public redo strands. But there is no way to tell if the redo for your change has been persisted yet, because that wait is gone too (that wait is the infamous ‘log file sync’ wait). If you want batching but still want a write notification, you must set commit_write to wait explicitly. By doing that you do not get the optimal batching because then waiting for the logwriter, including sending a signal to write is executed, which I suspect to be in the same ballpark as regular committing, but I haven’t checked that.

A look into oracle redo, part 9a: commit – concurrency considerations

April 9, 2018

Exadata, Linux, Oracle, Oracle EE, Oracle XE

1 Comment

During the investigations of my previous blogpost about what happens during a commit and when the data becomes available, I used breaks in gdb (GNU debugger) at various places of the execution of an insert and a commit to see what is visible for other sessions during the various stages of execution of the commit.

However, I did find something else, which is very logical, but is easily overlooked: at certain moments access to the table is blocked/serialised in order to let a session make changes to blocks belonging to the table, or peripheral blocks like undo, for the sake of consistency. These are changes made at the physical layer of an Oracle segment, the logical model of Oracle says that writers don’t block readers.

The first question is if this is a big issue. Well, for that you should simply look at any normal usage of an Oracle database. I don’t believe these serialisation moments are an issue, because the majority of my client’s does not experience serialisation problem.

However, this can become an issue if database processes can not run at will. What that means is that if processes are randomly stopped from execution by the operating system or hypervisor, it could be that the database process can be at an earlier mentioned ‘serialisation point’, which then means that access for all other processes remains blocked until the process is made running again, which then gives the process the opportunity to free exclusive access to a resource.

What are issues when a database process can not run at will?
– CPU oversubscription. If more processes are running than CPU’s are available (think about core’s and threads, and what it means for your CPU architecture), it means the operating system needs to make a decision on what it wants to be running, and what it wants to wait. No operating system scheduler understands or can know when an Oracle database process is running in a critical part of code, and therefore should not stop execution of it.
– Memory oversubscription. Whenever your system starts actively swapping in and out, your system is swapping, which means it’s moving active pages onto the swap device, and reads back active pages from the swap device for a process that demands them. This will activate additional tasks, and make processes get stuck waiting for memory to become available at all kinds of points during execution. Often, swapping is the beginning of the end, which means the services on the server stop functioning altogether for multiple reasons, like the Out Of Memory killer, or simply timing out for client’s of a process.
– Oversubscription of CPU or memory using virtualisation. As a troubleshooter, I would say this is actually worse than the ‘simple oversubscription’ mentioned in the two earlier points. Because typically, the visible part is the virtual machine. The virtual machine in this case itself is not oversubscribed, the oversubscription takes place at a layer invisible to the virtual machine, which unexplainably performs unpredictable. It is simply not okay to oversubscribe any machine that needs to run something that is latency sensitive like the Oracle database, unless you fully understand all the trade-offs that are coming with it, and you have the ability to understand when it does start to occur.

Whenever database processes can not run at will, you will see waits you didn’t see when processes could run at will. Typical waits in that case are (not an exhaustive list):
– any latch waits
– buffer busy waits
– any cursor: pin waits
Please mind these waits are not unique for this issue, any of these waits can occur for other reasons too.

Let me show a couple of examples for which the wait is artificially created by stopping execution at a point where this serialisation takes place, thereby mimicking getting put off cpu due to load, using the earlier insert and commit execution:

a) latch: cache buffers chains
Whenever a process needs to read or modify a database buffer, it needs to ‘pin’ the buffer. Such a pin is also known as a buffer handle or a buffer state object. The pin, which is a memory area, must be obtained from a freelist and attached to the buffer header in order to pin the buffer. A pin is obtained using the kcbzgs (kernel cache buffers helper functions get state object) function, which calls the kssadf_numa_intl (kernel service state objects add from freelist numa internal) function which initialises the state object.

The function that performs the actual pin of a block is not a single one, pinning a block is done in multiple functions. Two of them are kcbgtcr (kernel cache buffers get consistent read) and kcbgcur (kernel cache buffers get current). Under normal circumstances, concurrency for a consistent read of a buffer (kcbgtcr) does not lead to any blocking, because the latch (cache buffers chains) can be taken in shared mode, and the pinning is done in ‘fast mode’ (my interpretation):

> kcbgtcr(0x7f88a8b79db0, 0, ...)
| > kscn_to_ub8_impl(0x7f88a8b88d4c, 0x6baefdc0, ...)
| < kscn_to_ub8_impl+0x00000000003e returns: 0x287562
| > ktrexf(0x7fff998123f8, 0x7f88ae5d7f30, ...)
| | > ktrEvalBlockForCR(0x7f88a8b88d4c, 0x7f88ae5d7f30, ...)
| | < ktrEvalBlockForCR+0x0000000000f4 returns: 0x1
| < ktrexf+0x0000000000a4 returns: 0x9
| > kcbzgsf(0x1, 0x7fff998120b8, ...)
| | > kssadf_numa_intl(0x36, 0x6ddc0b50, ...)
| | | > kpdbIdToUid(0, 0x6ddc0b50, ...)
| | | < kpdbIdToUid+0x0000000000e7 returns: 0
| | < kssadf_numa_intl+0x000000000235 returns: 0x6bf84908
| < kcbzgsf+0x0000000001a0 returns: 0x6bf84908
| > kcbz_fp_cr(0x8dff9478, 0x6bf84998, ...)
| < kcbz_fp_cr+0x000000000089 returns: 0
| > kscn_to_ub8_impl(0x7f88a8b88d70, 0x1, ...)
| < kscn_to_ub8_impl+0x00000000003e returns: 0
< kcbgtcr+0x000000000b38 returns: 0x8df94014

Inside the kcbgtcr function, kcbzgsf (kernel cache buffers helper functions test state object fast) is called, and after that, the state object is pinned to the block in kcbz_fp_cr (kernel cache buffers helper functions fast pin for consistent read). Do you see the absence of ksl_get_shared_latch function to serialise access to add this pin? Actually the cache buffers chains latch is gotten, but it’s done inside the kcbgtcr function. The latch is gotten before the ktrexf (kernel transaction redo examine block fast) in shared mode, and the latch is released after the kcbz_fp_cr function. Multiple processes can pin the block for consistent read without serialisation or blocking each other because of the cache buffers chains latch being a shared latch, and this latch is taken in shared mode.

Why do I show this? Well, in the “old” days, which is Oracle 8.1, Oracle did not use shared latches (at least on the platforms I worked on, but if my memory serves me right, shared latches were introduced starting from Oracle 9), which meant for any logical read the cache buffers chains latch had to be obtained. This could get painful for very frequently visited blocks, like index root blocks, and high concurrency, and additional to that, because there was no post/wait messaging for latch waits but only spinning, this meant lots of processes waiting for latch number 98, and CPU usage going through the roof.

Back to today. Below is shown how the function calls look like when DML is done. When a buffer change is done (data is changed), a buffer has to be obtained in current mode, which is done using the kcbgcur function:

> kcbgcur(0x7ffda5957c70, 0x2, ...)
| > kcbzgs(0x1, 0x2, ...)
| | > kssadf_numa_intl(0x36, 0x7a6086c8, ...)
| | | > kpdbIdToUid(0, 0x7a6086c8, ...)
| | | < kpdbIdToUid+0x0000000000e7 returns: 0
| | < kssadf_numa_intl+0x000000000235 returns: 0x7867c9c0
| < kcbzgs+0x000000000178 returns: 0x7867c9c0
| > kti_is_imu(0x7867ca50, 0xb0f71018, ...)
| | > ktcgtx(0x76f29d50, 0xb0f71018, ...)
| | < ktcgtx+0x00000000001e returns: 0x76f29d50
| < kti_is_imu+0x000000000039 returns: 0
< kcbgcur+0x0000000009fb returns: 0xb0302014

This looks reasonably the same as the previous call overview which shown kcbgtcr, however the kcbzgsf function changed to kcbzgs, so minus the ‘f’ for fast, and a few other functions are missing. Another interesting thing (not visible) is the cache buffers chains latch is obtained after the kcbzgs function that obtains the buffer state object, and the latch is obtained with special bit 0x1000000000000000 set to indicate the shared latch is obtained in non-shared mode. After the kti_is_imu function the pin is attached to the buffer and the latch is freed.

To make things a little more complicated, a buffer can be gotten in current mode, but still get the cache buffers chains in shared mode. This is how that looks like:

> kcbgcur(0x7ffda5955620, 0x1, ...)
| > kcbzgs(0x1, 0x1, ...)
| | > kssadf_numa_intl(0x36, 0x7a6086c8, ...)
| | | > kpdbIdToUid(0, 0x7a6086c8, ...)
| | | < kpdbIdToUid+0x0000000000e7 returns: 0
| | < kssadf_numa_intl+0x000000000235 returns: 0x7867c9c0
| < kcbzgs+0x000000000178 returns: 0x7867c9c0
| > kcbz_fp_shr(0xb0fa3cb8, 0x7867ca50, ...)
| < kcbz_fp_shr+0x000000000085 returns: 0x7f3011d82f40
< kcbgcur+0x0000000009fb returns: 0xb07a0014

The function kcbzgs without ‘f’ is called, but the cache buffers chains latch is gotten in shared mode (not visible; done in the kcbgcur function after kcbzgs), and there’s the function kcbz_fp_shr (kernel cache buffers helper function fast pin shared) to pin the buffer in shared mode.

In order to be allowed to change a buffer, the cache buffers chains latch must be taken in non-shared mode to change the buffer header to indicate the buffer is busy and guarantee only the process itself can access it. Obtaining the cache buffers chains latch in non-shared mode means that access to the hash bucket (multiple hash buckets actually) to which the buffer is linked is blocked until the latch is freed. Please mind the period that the latch is obtained here is very short.

However, this blocking scenario can be easily replayed:
1. session 1: startup a database foreground session.
2. session 2: obtain the PID of the session 1’s process and attach to it with gdb.
3. session 2: break on the kti_is_imu function (break kti_is_imu) and continue.
4. session 1: insert a row into a table (insert into test values (‘a’);). the breakpoint will fire in gdb, stopping execution.
5. session 3: startup another database foreground session.
6. session 3: select the table on which the insert just broke (select * from test;). This will hang.
7. session 4: startup a SYSDBA session, and obtain the wait state of the foreground sessions:

SQL> @who active

SID_SERIAL_INST OSPID    USERNAME      C SERVER    PROGRAM                                          EXTRA
--------------- -------- ------------- - --------- ------------------------------------------------ -------------
59,53450,@1     28166    SYS           * DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:0ms,event:ON CPU:SQL,seq#:98
102,53150,@1    28083    TS              DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:66867.83ms,event:ON CPU:SQL,seq#:33
101,6637,@1     28186    TS              DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:344013.82ms,event:latch: cache buffers chainsp1:1811473056p2:289p3:24599> Blocked by (inst:sid): 1:102,seq#:39

SID 101 is waiting on the cache buffers chains latch, because SID 102 is holding that. SID 102 is showing ON CPU because we broke execution outside of a wait event, so Oracle thinks it’s busy executing, while it is actually stopped by gdb.

Again, this is a wait (latch: cache buffers chains) that is unlikely be an issue, and if it is, you either ran into a bug, or this is only the way a much bigger problem (oversubscription, bad application design) is showing itself.

b) buffer busy waits
The name of the wait event ‘buffer busy wait’ is self explanatory. The technical implementation is way lesser known. What technically happens during a buffer busy wait, is that a session needs to read a certain buffer, and in order to read it, it performs the usual actions of selecting the correct cache buffers chains bucket, following the double linked list to the buffer headers to find the best best buffer and then obtain a buffer state object. The critical part for running into the buffer busy wait is when the session checks the state of the buffer via the buffer header and finds the change state (X$BH.CSTATE) to be higher than 0, indicating the block is currently being changed. At this point the session calls kcbzwb (kernel cache buffers helper function wait for buffer), and must wait for the buffer state to be changed and therefore accessible, for which the wait event is the buffer busy wait event. When the change state is reverted back to 0, the buffer is accessible again.

Once a buffer is pinned in current mode using kcbgcur en kcbzgs, the buffer is still fully accessible for all sessions. Only once the database truly starts to change the buffer, which is done in kcbchg1_main, the cache buffers chains latch is taken non-shared, and the CSTATE field in the buffer header is changed to 1 and then the latch is freed. My current investigations show that the CSTATE is changed to 2 after the kcopcv and kcrfw_copy_cv functions, so when the redo is written to the public redo strand, and set to 4 after the changes have been written to the buffers, which is done in kcbapl. At the end of the kcbchg1_main function, when all the changes are done, the cache buffers chains latch is taken non-shared again, the CSTATE field is set to 0 and the latch is freed to enable access to the buffer again.

68801:kcbchg1_main+2536:0x000000006ba4e4d8(Variable Size(pgsz:2048k)|+182772952 shared pool|(child)latch#5383:cache buffers chains+0 shared pool|permanent memor,duration 1,cls perm+2204888 ):R:8:0/0()
68802:kcbchg1_main+2536:0x000000006ba4e4d8(Variable Size(pgsz:2048k)|+182772952 shared pool|(child)latch#5383:cache buffers chains+0 shared pool|permanent memor,duration 1,cls perm+2204888 ):W:8:0x1000000000000022/1152921504606847010()
..
68897:kcbchg1_main+3318:0x00000000967eddda(Variable Size(pgsz:2048k)|+901701082 buffer header|9668A000+66 ):W:1:0x1/1()
..
68939:kcbchg1_main+3715:0x000000006ba4e4d8(Variable Size(pgsz:2048k)|+182772952 shared pool|(child)latch#5383:cache buffers chains+0 shared pool|permanent memor,duration 1,cls perm+2204888 ):R:8:0x1000000000000022/1152921504606847010()
68940:kcbchg1_main+3715:0x000000006ba4e4d8(Variable Size(pgsz:2048k)|+182772952 shared pool|(child)latch#5383:cache buffers chains+0 shared pool|permanent memor,duration 1,cls perm+2204888 ):W:8:0/0()

Above you see a cache buffers chains latch being obtained non shared (special bit 0x1000000000000000 is set), a little later the buffer header is updated which sets the value 1 at offset 66. Once this happens, this is visible in the CSTATE column of X$BH. A little further the latch is freed.

Why am I showing all of this? Well, there is quite some work that is done while the buffer is in a changed state (CSTATE>0) and therefore inaccessible. In normal situations, this is not a problem because despite all the work this is done very quickly, and therefore it’s hardly noticeable/measurable, so there’s no or very little time spend in buffer busy waits in most databases

However… if processes do get stuck randomly, it’s possible that a process gets stuck while making changes to a buffer, which then means any additional access to the buffer results in buffer busy waits. Also mind that a change to a table buffer requires (at least; in the most simple case) three buffers to be changed: the buffer containing the table data, the buffer containing the undo segment transaction table and the buffer containing the actual undo. A commit requires one buffer: the undo segment transaction table.

Let me show you an example:
1. session 1: startup a database foreground session.
2. session 2: obtain the PID of the session 1’s process and attach to it with gdb.
3. session 2: break on ktugur and kcopcv and continue.
4. session 3: startup a database foreground session.
5. session 1: insert a row into a table: insert into test values (‘a’);
6. session 2: gdb breaks execution of session 1 because it encountered ktugur.
7. session 3: select * from test; this runs unblocked.
8. session 4: startup a database foreground session as sysdba.
9. session 4: select * from v$lock; this will hang.

This one wasn’t obvious. I found this upon doing the tests. When a foreground session is in ktugur, it is creating the change vectors (kernel transaction undo generate undo and redo), and it holds two buffers in current mode:

SQL> l
  1  select 	ts.name tablespace,
  2  	dbablk,
  3  	class,
  4  	state,
  5  	decode(state,0,'free',1,'xcur',2,'scur',3,'cr', 4,'read',5,'mrec',6,'irec',7,'write',8,'pi', 9,'memory',10,'mwrite',11,'donated', 12,'protected',  13,'securefile', 14,'siop',15,'recckpt', 16, 'flashfree',  17, 'flashcur', 18, 'flashna') state_decoded,
  6  	mode_held,
  7  	changes,
  8  	cstate,
  9  	flag,
 10  	flag2,
 11  	rflag,
 12  	sflag,
 13  	lru_flag,
 14  	dirty_queue
 15  from 	  x$bh b
 16  	, v$tablespace ts
 17  where 	1=1
 18  and	not (us_nxt = us_prv and wa_nxt = wa_prv and to_number(wa_nxt,'xxxxxxxxxxxxxxxx') = to_number(us_nxt,'xxxxxxxxxxxxxxxx') + 16)
 19* and	b.ts#=ts.ts#
SQL> /

TABLESPACE                         DBABLK      CLASS      STATE STATE_DECO  MODE_HELD    CHANGES     CSTATE       FLAG      FLAG2      RFLAG      SFLAG   LRU_FLAG DIRTY_QUEUE
------------------------------ ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- -----------
UNDOTBS1                              128         17          1 xcur                2          1          0    2097152          0          0          0          0           0
TS                                  24597          1          1 xcur                2          1          0    2097153          0          0          0          0           0

The thing that was not obvious at first was that the blocks have CSTATE 0; so they can be accessed and read for a consistent read. Then I look at the wait and the p1/2/3 of the wait:

SID_SERIAL_INST OSPID    USERNAME      C SERVER    PROGRAM                                          EXTRA
--------------- -------- ------------- - --------- ------------------------------------------------ -------------------------------------------------------
100,262,@1      6274     SYS             DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:2519.48ms,event:buffer busy waitsp1:3p2:128p3:17> Blocked by (inst:sid): 1:59,seq#:476
142,15103,@1    6695     TS              DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:47115.23ms,event:SQL*Net message from client,seq#:226
12,25165,@1     7036     SYS           * DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:0ms,event:ON CPU:SQL,seq#:3165
59,31601,@1     5891     TS              DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:31292.51ms,event:ON CPU:SQL,seq#:111

Above SID 100 is session 4; it hangs on a buffer busy wait. The datafile is #2, which is the undo table space UNDOTBS1 and the buffer that it is waiting for is 128. Why does SID 100 needs to wait for the buffer, while another session (session 3 in the above runbook) can run without getting blocked for a buffer that is held exactly in the same state (xcur and CSTATE is zero)?

The answer can be gotten when executing pstack on the waiting process (or attach with gdb and obtain a backtrace):

(gdb) bt
#0  0x00007fd6a085bbda in semtimedop () at ../sysdeps/unix/syscall-template.S:81
#1  0x0000000010f9cca6 in sskgpwwait ()
#2  0x0000000010f9a2e8 in skgpwwait ()
#3  0x0000000010a66995 in ksliwat ()
#4  0x0000000010a65d25 in kslwaitctx ()
#5  0x00000000015c5e5e in kcbzwb ()
#6  0x0000000010b3748e in kcbgcur ()
#7  0x0000000010ab9b17 in ktuGetUsegHdr ()
#8  0x00000000014075d3 in ktcxbcProcessInPDB ()
#9  0x0000000001406de2 in ktcxbc ()
#10 0x0000000010ea8781 in qerfxFetch ()
#11 0x0000000010bb919c in rwsfcd ()
#12 0x00000000035ae618 in qeruaFetch ()
#13 0x00000000035ad837 in qervwFetch ()
#14 0x0000000010bb919c in rwsfcd ()
#15 0x0000000010ea1528 in qerhnFetch ()
#16 0x0000000010bb919c in rwsfcd ()
#17 0x0000000010ea1528 in qerhnFetch ()
#18 0x0000000010c9eeb3 in opifch2 ()
#19 0x000000000268c3de in kpoal8 ()
#20 0x0000000010ca5b3d in opiodr ()
#21 0x0000000010f37a29 in ttcpip ()
#22 0x0000000001eb4674 in opitsk ()
#23 0x0000000001eb914d in opiino ()
#24 0x0000000010ca5b3d in opiodr ()
#25 0x0000000001eb04ed in opidrv ()
#26 0x0000000002afedf1 in sou2o ()
#27 0x0000000000d05577 in opimai_real ()
#28 0x0000000002b09b31 in ssthrdmain ()
#29 0x0000000000d05386 in main ()

The important bit is at #6: kcbgcur (kernel cache buffers get buffer in current mode)! In other words: the block is gotten in current mode while another session is holding the block in current mode, which means it has to wait. The function kcbzwb means kernel cache buffer helper function wait for buffer.

I really wondered why the undo segment buffer is gotten in current mode; this is a read, there is no intention to change something. The current educated guess is this a way to be sure this buffer contains the current state of the transactions across all nodes in a RAC cluster. So this might be a clever optimisation to be absolutely sure the current state of the undo segment transaction table is gotten.

10. session 2: continue. This will break on kcopcv.
11. session 3: select * from test; this will hang.

Now we have two hanging sessions waiting for the wait event buffer busy wait:

SQL> @who active

SID_SERIAL_INST OSPID    USERNAME      C SERVER    PROGRAM                                          EXTRA
--------------- -------- ------------- - --------- ------------------------------------------------ -----------------
100,262,@1      6274     SYS             DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:120790ms,event:buffer busy waitsp1:3p2:128p3:35> Blocked by (inst:sid): 1:59,seq#:501
142,15103,@1    6695     TS              DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:1175715.03ms,event:buffer busy waitsp1:4p2:24597p3:1> Blocked by (inst:sid): 1:59,seq#:250
12,25165,@1     7036     SYS           * DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:0ms,event:ON CPU:SQL,seq#:3309
59,31601,@1     5891     TS              DEDICATED sqlplus@memory-presentation.local (TNS V1-V3)    time:9544.97ms,event:ON CPU:SQL,seq#:116

SID 100 still hangs in wait event buffer busy wait for file 3, block 128, but now the session doing the full table scan also hangs in wait event buffer busy wait, but for file 4, block 24597. Let’s look at the buffers held in current mode again:

SQL> @pinned_blocks

TABLESPACE                         DBABLK      CLASS      STATE STATE_DECO  MODE_HELD    CHANGES     CSTATE       FLAG      FLAG2      RFLAG      SFLAG   LRU_FLAG DIRTY_QUEUE
------------------------------ ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- ---------- -----------
UNDOTBS1                              128         35          1 xcur                2          0          1    2097156          0          0          0          0           0
UNDOTBS1                             3127         36          1 xcur                2          0          1    2097156          0          0          0          0           0
TS                                  24597          1          1 xcur                2          0          1    2097157          0          0          0          0           0

Because the insert progressed to kcopcv, the CSTATE of all three buffers involved in the insert, the undo segment transaction table, the undo segment undo entry and the buffer which has the inserted row are set to 1. This means that the two buffer busy waits are actually two different reasons: the select * from v$lock still waits to get a buffer in current mode, and the select * from test, which does a consistent read, now is waiting for the CSTATE to get set back to 0. In order to solve the hanging, go to session 2 and continue. This will break again on kcoapl, and again if you continue, because there are 3 blocks involved in this insert. Another way to make the insert finish is to disable all breakpoints (disable command) or to quit gdb.

Just for the fun of it, here is another example of how a stuck session can escalate: a process that gets stuck executing the kcbapl (kernel cache buffers apply a change to a buffer). This scenario can be replayed using the following steps:
1. session 1: startup a database foreground session.
2. session 2: obtain the PID of the session 1’s process and attach to it with gdb.
3. session 2: break on kcbapl and continue.
4. session 3: startup a database foreground session.
5. session 3: insert a row into a table: insert into test values (‘a’);
6. session 1: insert a row into a table: insert into test values (‘a’);
7. session 2: gdb broken execution of session 1 because it encountered kcbapl.
8. session 3: commit; this will hang

Can you reason why session 3 is hanging? Both the sessions are doing inserts to the same table…Give it a thought and replay it yourself.

Explanation: the breakpoint stopped execution before actually doing the changes. CSTATE for the buffer is set to 1, it will get set to 2 inside kcbapl to indicate the change vectors have been copied. But the important bit is that a redo copy latch was taken before executing kcbapl. The commit in session 3 sent a message to the logwriter asynchronously asking it to write (first invocation of kcrf_commit_force_int), and then freed the logical locking of the insert, and then invoked kcrf_commit_force again to wait for the logwriter to finish writing and get notified. However, the logwriter needs to obtain all the redo copy latches (it actually reads a metadata structure that indicates the redo copy latches are freed or not), at which it needs to wait because session 1 stuck holding a redo copy latch and therefore needs to wait. Not only does this show how dangerous it is when sessions can’t run, it also shows how dangerous it can be to run gdb in a production database, it can be quite hard to figure out what will be the implication of attaching with gdb to the process. So the actual outcome of above is that session 3 waits in the event ‘log file sync’, and the logwriter is waiting in ‘LGWR wait for redo copy’.

Summary
The Oracle database is carefully setup to be able to handle concurrency. In most normal situations you will see no to maybe very little time spend in wait events that indicate concurrency, like latches or buffer busy waits. If an Oracle database is on a machine that has more active processes than available CPU threads or cores (this depends on the hardware and CPU architecture; for Intel based architectures you should lean towards core’s, for SPARC you can lean more towards threads), the operating system needs to make choices, and will do that trying to give all processes an equal share. This means that processes will be put off CPU in order to give all processes a fair share of CPU run time. Also mind other oversubscription possibilities, like using more memory than physically available, and doing the same using virtualisation, which is handing out more VCPU’s than CPU threads, and handing out more memory than physically available. Because you can’t look outside your virtual machine, it’s often harder to troubleshoot virtualisation oversubscription issues.

Oracle really tries to minimise the amount of time spend while holding a non shared latch. Therefore with modern Oracle databases it’s unlikely that latch waits show up, even if a system is (not too much) oversubscribed. If you do have a latching problem, it’s most likely you have got another problem which causes this.

The most well documented reason for buffer busy waits is a buffer being written onto persistent media, and therefore the buffer is marked busy for the time of the write. Another reason that is documented, is getting a buffer in a mode that is not compatible with the mode it is held in. Having a buffer in current mode will block another request for the buffer in current mode. But, having a buffer pinned in current mode does not block consistent read access, it’s only when the buffer is being changed that requires the buffer to be marked busy. This state is externalised in the CSTATE column of the view X$BH, which shows fields of the buffer header struct that is used to manage the buffer cache buffers. Because there is quite some work to be done when a buffer is changed, and during making these changes the buffers are not usable (‘busy’) for other processes, a buffer busy wait is more likely to occur when oversubscription is taking place.

Thanks to Jonathan Lewis for explaining how to get the buffer header addresses.
Thanks to Abel Macias for brainstorming.

A look into oracle redo, part 9: commit

April 3, 2018

Exadata, Linux, Oracle, Oracle EE, Oracle XE

2 Comments

The previous blogpost talked about a simple insert, this blogpost investigates what happens when the DML is committed. Of course this is done with regular commit settings, which means means they are not touched, which means commit_logging is set to immediate and commit_wait is set to wait as far as I know. The documentation says there is no default value, and the settings are empty in all parameter views. In my humble opinion, if you must change the commit settings in order to make your application perform usable with the database, something is severely wrong somewhere.

This blogpost works best if you thoroughly gone through the previous post. I admit it’s a bit dry and theoretical, but you will appreciate the knowledge which you gained there, because it directly applies to a commit.

First let’s look at the flow of functions for the commit:

kpoal8
  opiexe
    kksExecuteCommand
      xctCommitTxn
        xctctl
          k2send
            k2lcom
              ktdcmt
                ktcCommitTxn
                  ktcCommitTxn_new
                    qesrcCM_Enabled
                    qesrcCM_PreCmt_
                      qesrcTol_New
                    ktuIMTabPresetTxn
                    ktucmt
                      kcbchg1                 
                        kcbchg1_main
                          ktiimu_chg
                          kcopcv
                          kcrfw_redo_gen_ext
                            kcrfw_copy_cv
                            kcscur_rba
                            kcbapl
                              kcbhlchk
                              kcoapl
                              kco_issue_callback
                                kturcm
                              kco_blkchk
                              kcoadv_hdr
                              kcbhfix_tail
                            kcrfw_partial_checksum
                            kcrf_commit_force_int
                              kcscu8
                              kcscu8
                              ksbasend
                          ktuulc
                            ksqrcl
                              ksqrcli_int
                    ktudnx   
                  kssdct
                    kssdch_int
                      kssdel
                        kss_del_cb
                          ktldbl
                            ktldbl_noredo
                              kcbnlc
                                ktbdbc
                    kssdch_int
                      kssdel
                        kss_del_cb
                          ktaidm
                            ksqrcl
ksupop
  kcrf_commit_force_int
    kcscu8
    kcscu8
    ksbasend
    kslawe
    kcscu8
    kslwtbctx
    kslwaitctx
  - kcscu8
    kslawe
    kcscu8
    kslwaitctx
  - kcscu8
    kslawe
    kcscu8
    kslwaitctx
  - kcscu8
    kslwtectx

Please mind this is an overview of functions which is not complete, it provides enough information to show the flow of functions I want to highlight. There are much more functions involved during the execution of commit.

The first thing that is visible here is that after the kpoal8/opiexe/kksExecuteCommand (kernel compile shared objects Execute Command) function are xct (transaction control) functions. Of course this is logical, a commit ends a transaction and makes changes visible. The xct layer then moves into the k2 layer, which is the distributed execution layer. I am not doing anything distributed, it is my current understanding that this layer is invoked this way so that if anything distributed was pending, it would be handled appropriately. After the k2 layer the function ktdcmt (kernel transaction distributed commit) is executed.

After the distributed layers we enter the ktc layer (kernel transaction control). In ktcCommitTXN_new I see handling of features like the result cache (qesrc) and In-Memory (ktuIM), then the ktu layer (kernel transaction undo) is entered, which enters the kcb layer (kernel cache buffers) using functions we saw in the previous post: kcbchg1 and kcbchg1_main.

In fact, at this point it looks very similar to the insert, in kcbchg1_main, ktiimu_chg and kcopcv (prepare change vectors) are called, but only once (because a commit only involves one block, see a little further for the explanation) instead of three times as we saw with the insert. Then kcrfw_redo_gen_ext is called, which is doing almost the same as the insert: first kcrfw_copy_cv is executed to copy the change vector to the public redo strand, then kcbapl is called to apply the change to a buffer. The kco_issue_callback function calls kturcm (kernel transaction undo redo commit) indicating the type change to the buffer. This means that a commit changes a single buffer, which is the buffer that holds the transaction in the transaction table in the undo segment, and the change is marking the transaction as committed. So a ‘commit marker’ is not a special token that is written into the redo stream, but in fact it’s simply a block change, just like all other change vectors.

After kcbapl, kcrfw_partial_checksum is called to checksum the redo in the public redo strand, again just like we saw with the insert.

Unique to a commit is the invocation of the kcrf_commit_force_int function. This is the first ‘use’ of the kcrf_commit_force_int function (indicated by the second function argument set to zero, not visible in the overview), which is signalling the logwriter to write any unwritten change vectors in the public redo strands. kcrf_commit_force_int checks the on disk SCN and the LWN SCN using kcscu8 (kernel cache scn management read current SCN) in the kcrfsg_ struct to check logwriter progress:
– If the on disk SCN is beyond the process’ commit SCN, the change vectors are written, and no further action is necessary (this function is quit), which also means a second invocation of kcrf_commit_force_int is not necessary.
– If the on disk SCN isn’t progressed beyond the process’ commit SCN, but the LWN SCN is, it means the logwriter is currently writing the change vectors for this commit SCN. In that case there is no need to signal the logwriter, but it requires the process to validate the write later using the second invocation of kcrf_commit_force_int.
– If both the on disk SCN and LWN SCN did not progress beyond or are equal to the process’ commit SCN, this invocation of kcrf_commit_force_int needs to send the logwriter a message using ksbasend (kernel service background processes asynchronous send message) to start writing the public redo strands. ksbasend will only send a message if the messages flag in the kcrfsg_ struct is not set indicating it has already been signalled.
After which the kcrf_commit_force_int function is returned from, as well as the kcrfw_redo_gen_ext function, so we are back in kcbchg1_main.

Also different from an insert is the invocation of the ktuulc (kernel transaction undo unlock; this is a guess) function. Which calls ksqrcl (kernel service enqueue release an enqueue), which calls ksqrcli_int (my guess this (=the addition of _int) is an enhanced version of the enqueue release function), which performs the clearance of the TX enqueue set for the inserted row. This clearance is not a guess, ksqrcli_int does clear the TX enqueue for the inserted row. After clearing the row lock, some more functions returned from: kcbchg1_main and kcbchg1, so we are back in ktucmt.

Because the transaction is now committed, the active transaction count in the undo segment can be decreased, which is done in ktudnx (kernel transaction undo decrease active transaction count). Then the ktucmt function is returned from too, and we are back in ktcCommitTxn_new.

In ktcCommitTxn_new state objects are cleaned up using kssdct (kernel service state object delete children of specified type). A state object is a memory area that keeps the state of various things that are transient, so if they get lost, the state object reflects the last known state. The callback action of the function performs some more housekeeping, the ktldbl (kernel transaction list blocks changed delete list of block SO’s) function removes block SO’s/buffer handles, which calls kcbnlc (kernel cache buffers analyse cleanout), which calls ktbdbc (kernel transaction block fast block cleanout) to perform delayed block cleanout/commit cleanout. This cleans up the state in the data block, which means it cleans up the lock bytes, set Fsc/Scn to the commit SCN and set the commit flag to C— in the ITL in the block.

The next state object that is cleaned is the shared table lock (TM); by looking at the functions it’s quite easy to understand that this is happening, ksqrcl is the function to release and enqueue, and ktaidm is kernel transaction access internal deallocate dml lock function.

Past releasing the TM enqueue, there are other things done for which I didn’t put their function names in, but the execution is returning from a lot of the other functions shown as calling functions. Of course Oracle needs to update all kind of counters and usage statistics, and record audit information. But eventually, everything has been released. However, there is something more that is executed as part of a commit. This is the second invocation of the kcrf_commit_force_int function.

Actually, when kcrf_commit_force_int is executed for the second ‘use’, this is visible with the second argument of the calling arguments is set to ‘1’ (not visible in the function call overview above). The functions that are executed in kcrf_commit_force_int are actually exactly the same as the first invocation:
– kcscu8 is called to read the on disk SCN from kcrfsg_
– kcscu8 is called to read the LWN SCN from kcrfsg_
Also the same logic is applied to the values that are the result of calling kcscu8 to read the SCN values as stated previously. If these SCNs did not progress far enough, ksbasend is called.

The interesting thing of the second execution of kcrf_commit_force_int happens after ksbasend: the kcrf_commit_force_int function loops until the on disk SCN has progressed beyond the process’ commit SCN (which means the change vectors are written from the public redo strands into the online redologfiles). To indicate it’s waiting/looping for the on disk SCN to progress that far, the wait interface is called (kslwtbctx) for the wait ‘log file sync’, after which it loops, for which I put a hyphen before the start of the loop to indicate what to loop consists of.

I illustrated the post/wait mode of log file sync, which is visible with ‘kslawe’ (kernel service lock management add post-wait entry). The post-wait entry is removed inside kslwaitctx, and then setup again. Interestingly, when in post-wait mode, the process must be posted by the logwriter, even if it finds the on disk SCN to have progressed beyond the process’ commit SCN. The other mode for waiting for the on disk SCN is called ‘polling’, search my blog for articles about it if this sparked your interest.

Summary
The intention of this blogpost is not to bury you in Oracle internal functions, despite the look of the article and the amount of functions mentioned :-). The aim for spelling out the functions is to show what happens, and to learn about the layers in which they execute.

If you skip past the first couple of functions that are executed with a commit, the ktc (kernel transaction control) layer is crossed, then the ktu (kernel transaction undo) layer, after which the change is executed under supervision of the kcb (kernel cache buffer) layer.

In fact, the rough outline of the change is the same as described in the previous article about insert: kcbchg1, kcbchg1_main, kcopcv, kcrfw_redo_gen_ext, etc. Just like with the insert, the function called in the function kco_issue_callback sets the type of block change, which is kturcm with commit.

A commit is a change to the block that holds the undo segment’s transaction table, and flags the current transaction as committed. This is what is commonly referred to as a ‘commit marker’.

After the kturcm function, the transaction is changed to the status committed. However, if you look closely, there are several functions executed AFTER kturcm, like kco_blkchk, kcoadv_hdr and kcbhfix_tail that complete the change made in kturcm in order to make the block consistent.

After block changes and the change vector checksum in kcrfw_partial_checksum, a function unique to commit is executed: kcrf_commit_force_int. The first time invocation of this function signals the logwriter to write.

At the time of kcrf_commit_force_int and returning from it into the function kcrfw_redo_gen_ext back to kcbchg1_main, the newly inserted value is not available, but when the execution in the kcbchg1_main function reaches ktuulc to clean up the TX enqueue, the the NEW value becomes available!

This is something which I still do find counter intuitive, because this means at the above mentioned time, which is prior to reaching ktuulc the change becomes visible to all sessions but the committing session. The committing session at that point needs to clean up the block a little, and later on remove the shared TM enqueue, and after that, the committing session executes kcrf_commit_force_int again to wait for the ‘commit marker’ and obviously all successive change vectors to complete. WHILE ALL OTHER SESSIONS CAN SEE AND USE THE CHANGED DATA FOR WHICH THE COMMITTING SESSION IS WAITING!

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A look into oracle redo, part 11: log writer worker processes

A look into oracle redo: index and overview

A look into oracle redo, part 10: commit_wait and commit_logging

A look into oracle redo, part 9a: commit – concurrency considerations

A look into oracle redo, part 9: commit

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