<!-- $PostgreSQL: pgsql/doc/src/sgml/perform.sgml,v 1.84 2010/07/26 20:28:49 petere Exp $ -->

 <chapter id="performance-tips">
  <title>Performance Tips</title>

  <indexterm zone="performance-tips">
   <primary>performance</primary>
  </indexterm>

  <para>
   Query performance can be affected by many things. Some of these can
   be controlled by the user, while others are fundamental to the underlying
   design of the system.  This chapter provides some hints about understanding
   and tuning <productname>PostgreSQL</productname> performance.
  </para>

  <sect1 id="using-explain">
   <title>Using <command>EXPLAIN</command></title>

   <indexterm zone="using-explain">
    <primary>EXPLAIN</primary>
   </indexterm>

   <indexterm zone="using-explain">
    <primary>query plan</primary>
   </indexterm>

   <para>
    <productname>PostgreSQL</productname> devises a <firstterm>query
    plan</firstterm> for each query it receives.  Choosing the right
    plan to match the query structure and the properties of the data
    is absolutely critical for good performance, so the system includes
    a complex <firstterm>planner</> that tries to choose good plans.
    You can use the
    <xref linkend="sql-explain"> command
    to see what query plan the planner creates for any query.
    Plan-reading is an art that deserves an extensive tutorial, which
    this is not; but here is some basic information.
   </para>

   <para>
    The structure of a query plan is a tree of <firstterm>plan nodes</>.
    Nodes at the bottom level of the tree are table scan nodes: they return raw rows
    from a table.  There are different types of scan nodes for different
    table access methods: sequential scans, index scans, and bitmap index
    scans.  If the query requires joining, aggregation, sorting, or other
    operations on the raw rows, then there will be additional nodes
    above the scan nodes to perform these operations.  Again,
    there is usually more than one possible way to do these operations,
    so different node types can appear here too.  The output
    of <command>EXPLAIN</command> has one line for each node in the plan
    tree, showing the basic node type plus the cost estimates that the planner
    made for the execution of that plan node.  The first line (topmost node)
    has the estimated total execution cost for the plan; it is this number
    that the planner seeks to minimize.
   </para>

   <para>
    Here is a trivial example, just to show what the output looks like:
    <footnote>
     <para>
      Examples in this section are drawn from the regression test database
      after doing a <command>VACUUM ANALYZE</>, using 8.2 development sources.
      You should be able to get similar results if you try the examples yourself,
      but your estimated costs and row counts might vary slightly
      because <command>ANALYZE</>'s statistics are random samples rather
      than exact.
     </para>
    </footnote>

<programlisting>
EXPLAIN SELECT * FROM tenk1;

                         QUERY PLAN
-------------------------------------------------------------
 Seq Scan on tenk1  (cost=0.00..458.00 rows=10000 width=244)
</programlisting>
   </para>

   <para>
    The numbers that are quoted by <command>EXPLAIN</command> are (left
    to right):

    <itemizedlist>
     <listitem>
      <para>
       Estimated start-up cost (time expended before the output scan can start,
       e.g., time to do the sorting in a sort node)
      </para>
     </listitem>

     <listitem>
      <para>
       Estimated total cost (if all rows are retrieved, though they might
       not be; e.g., a query with a <literal>LIMIT</> clause will stop
       short of paying the total cost of the <literal>Limit</> plan node's
       input node)
      </para>
     </listitem>

     <listitem>
      <para>
       Estimated number of rows output by this plan node (again, only if
       executed to completion)
      </para>
     </listitem>

     <listitem>
      <para>
       Estimated average width (in bytes) of rows output by this plan
       node
      </para>
     </listitem>
    </itemizedlist>
   </para>

   <para>
    The costs are measured in arbitrary units determined by the planner's
    cost parameters (see <xref linkend="runtime-config-query-constants">).
    Traditional practice is to measure the costs in units of disk page
    fetches; that is, <xref linkend="guc-seq-page-cost"> is conventionally
    set to <literal>1.0</> and the other cost parameters are set relative
    to that.  (The examples in this section are run with the default cost
    parameters.)
   </para>

   <para>
    It's important to note that the cost of an upper-level node includes
    the cost of all its child nodes.  It's also important to realize that
    the cost only reflects things that the planner cares about.
    In particular, the cost does not consider the time spent transmitting
    result rows to the client, which could be an important
    factor in the real elapsed time; but the planner ignores it because
    it cannot change it by altering the plan.  (Every correct plan will
    output the same row set, we trust.)
   </para>

   <para>
    The <literal>rows</> value is a little tricky
    because it is <emphasis>not</emphasis> the
    number of rows processed or scanned by the plan node.  It is usually less,
    reflecting the estimated selectivity of any <literal>WHERE</>-clause
    conditions that are being
    applied at the node.  Ideally the top-level rows estimate will
    approximate the number of rows actually returned, updated, or deleted
    by the query.
   </para>

   <para>
    Returning to our example:

<programlisting>
EXPLAIN SELECT * FROM tenk1;

                         QUERY PLAN
-------------------------------------------------------------
 Seq Scan on tenk1  (cost=0.00..458.00 rows=10000 width=244)
</programlisting>
   </para>

   <para>
    This is about as straightforward as it gets.  If you do:

<programlisting>
SELECT relpages, reltuples FROM pg_class WHERE relname = 'tenk1';
</programlisting>

    you will find that <classname>tenk1</classname> has 358 disk
    pages and 10000 rows.  The estimated cost is computed as (disk pages read *
    <xref linkend="guc-seq-page-cost">) + (rows scanned *
    <xref linkend="guc-cpu-tuple-cost">).  By default,
    <varname>seq_page_cost</> is 1.0 and <varname>cpu_tuple_cost</> is 0.01,
    so the estimated cost is (358 * 1.0) + (10000 * 0.01) = 458.
   </para>

   <para>
    Now let's modify the original query to add a <literal>WHERE</> condition:

<programlisting>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 7000;

                         QUERY PLAN
------------------------------------------------------------
 Seq Scan on tenk1  (cost=0.00..483.00 rows=7033 width=244)
   Filter: (unique1 &lt; 7000)
</programlisting>

    Notice that the <command>EXPLAIN</> output shows the <literal>WHERE</>
    clause being applied as a <quote>filter</> condition; this means that
    the plan node checks the condition for each row it scans, and outputs
    only the ones that pass the condition.
    The estimate of output rows has been reduced because of the <literal>WHERE</>
    clause.
    However, the scan will still have to visit all 10000 rows, so the cost
    hasn't decreased; in fact it has gone up a bit (by 10000 * <xref
    linkend="guc-cpu-operator-cost">, to be exact) to reflect the extra CPU
    time spent checking the <literal>WHERE</> condition.
   </para>

   <para>
    The actual number of rows this query would select is 7000, but the <literal>rows</>
    estimate is only approximate.  If you try to duplicate this experiment,
    you will probably get a slightly different estimate; moreover, it will
    change after each <command>ANALYZE</command> command, because the
    statistics produced by <command>ANALYZE</command> are taken from a
    randomized sample of the table.
   </para>

   <para>
    Now, let's make the condition more restrictive:

<programlisting>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 100;

                                  QUERY PLAN
------------------------------------------------------------------------------
 Bitmap Heap Scan on tenk1  (cost=2.37..232.35 rows=106 width=244)
   Recheck Cond: (unique1 &lt; 100)
   ->  Bitmap Index Scan on tenk1_unique1  (cost=0.00..2.37 rows=106 width=0)
         Index Cond: (unique1 &lt; 100)
</programlisting>

    Here the planner has decided to use a two-step plan: the bottom plan
    node visits an index to find the locations of rows matching the index
    condition, and then the upper plan node actually fetches those rows
    from the table itself.  Fetching the rows separately is much more
    expensive than sequentially reading them, but because not all the pages
    of the table have to be visited, this is still cheaper than a sequential
    scan.  (The reason for using two plan levels is that the upper plan
    node sorts the row locations identified by the index into physical order
    before reading them, to minimize the cost of separate fetches.
    The <quote>bitmap</> mentioned in the node names is the mechanism that
    does the sorting.)
   </para>

   <para>
    If the <literal>WHERE</> condition is selective enough, the planner might
    switch to a <quote>simple</> index scan plan:

<programlisting>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 3;

                                  QUERY PLAN
------------------------------------------------------------------------------
 Index Scan using tenk1_unique1 on tenk1  (cost=0.00..10.00 rows=2 width=244)
   Index Cond: (unique1 &lt; 3)
</programlisting>

    In this case the table rows are fetched in index order, which makes them
    even more expensive to read, but there are so few that the extra cost
    of sorting the row locations is not worth it.  You'll most often see
    this plan type for queries that fetch just a single row, and for queries
    that have an <literal>ORDER BY</> condition that matches the index
    order.
   </para>

   <para>
    Add another condition to the <literal>WHERE</> clause:

<programlisting>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 3 AND stringu1 = 'xxx';

                                  QUERY PLAN
------------------------------------------------------------------------------
 Index Scan using tenk1_unique1 on tenk1  (cost=0.00..10.01 rows=1 width=244)
   Index Cond: (unique1 &lt; 3)
   Filter: (stringu1 = 'xxx'::name)
</programlisting>

    The added condition <literal>stringu1 = 'xxx'</literal> reduces the
    output-rows estimate, but not the cost because we still have to visit the
    same set of rows.  Notice that the <literal>stringu1</> clause
    cannot be applied as an index condition (since this index is only on
    the <literal>unique1</> column).  Instead it is applied as a filter on
    the rows retrieved by the index.  Thus the cost has actually gone up
    slightly to reflect this extra checking.
   </para>

   <para>
    If there are indexes on several columns referenced in <literal>WHERE</>, the
    planner might choose to use an AND or OR combination of the indexes:

<programlisting>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 100 AND unique2 &gt; 9000;

                                     QUERY PLAN
-------------------------------------------------------------------------------------
 Bitmap Heap Scan on tenk1  (cost=11.27..49.11 rows=11 width=244)
   Recheck Cond: ((unique1 &lt; 100) AND (unique2 &gt; 9000))
   -&gt;  BitmapAnd  (cost=11.27..11.27 rows=11 width=0)
         -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..2.37 rows=106 width=0)
               Index Cond: (unique1 &lt; 100)
         -&gt;  Bitmap Index Scan on tenk1_unique2  (cost=0.00..8.65 rows=1042 width=0)
               Index Cond: (unique2 &gt; 9000)
</programlisting>

    But this requires visiting both indexes, so it's not necessarily a win
    compared to using just one index and treating the other condition as
    a filter.  If you vary the ranges involved you'll see the plan change
    accordingly.
   </para>

   <para>
    Let's try joining two tables, using the columns we have been discussing:

<programlisting>
EXPLAIN SELECT *
FROM tenk1 t1, tenk2 t2
WHERE t1.unique1 &lt; 100 AND t1.unique2 = t2.unique2;

                                      QUERY PLAN
--------------------------------------------------------------------------------------
 Nested Loop  (cost=2.37..553.11 rows=106 width=488)
   -&gt;  Bitmap Heap Scan on tenk1 t1  (cost=2.37..232.35 rows=106 width=244)
         Recheck Cond: (unique1 &lt; 100)
         -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..2.37 rows=106 width=0)
               Index Cond: (unique1 &lt; 100)
   -&gt;  Index Scan using tenk2_unique2 on tenk2 t2  (cost=0.00..3.01 rows=1 width=244)
         Index Cond: (unique2 = t1.unique2)
</programlisting>
   </para>

   <para>
    In this nested-loop join, the outer (upper) scan is the same bitmap index scan we
    saw earlier, and so its cost and row count are the same because we are
    applying the <literal>WHERE</> clause <literal>unique1 &lt; 100</literal>
    at that node.
    The <literal>t1.unique2 = t2.unique2</literal> clause is not relevant yet,
    so it doesn't affect the row count of the outer scan.  For the inner (lower) scan, the
    <literal>unique2</> value of the current outer-scan row is plugged into
    the inner index scan to produce an index condition like
    <literal>unique2 = <replaceable>constant</replaceable></literal>.
    So we get the same inner-scan plan and costs that we'd get from, say,
    <literal>EXPLAIN SELECT * FROM tenk2 WHERE unique2 = 42</literal>.  The
    costs of the loop node are then set on the basis of the cost of the outer
    scan, plus one repetition of the inner scan for each outer row (106 * 3.01,
    here), plus a little CPU time for join processing.
   </para>

   <para>
    In this example the join's output row count is the same as the product
    of the two scans' row counts, but that's not true in all cases because
    you can have <literal>WHERE</> clauses that mention both tables
    and so can only be applied at the join point, not to either input scan.
    For example, if we added
    <literal>WHERE ... AND t1.hundred &lt; t2.hundred</literal>,
    that would decrease the output row count of the join node, but not change
    either input scan.
   </para>

   <para>
    One way to look at variant plans is to force the planner to disregard
    whatever strategy it thought was the cheapest, using the enable/disable
    flags described in <xref linkend="runtime-config-query-enable">.
    (This is a crude tool, but useful.  See
    also <xref linkend="explicit-joins">.)

<programlisting>
SET enable_nestloop = off;
EXPLAIN SELECT *
FROM tenk1 t1, tenk2 t2
WHERE t1.unique1 &lt; 100 AND t1.unique2 = t2.unique2;

                                        QUERY PLAN
------------------------------------------------------------------------------------------
 Hash Join  (cost=232.61..741.67 rows=106 width=488)
   Hash Cond: (t2.unique2 = t1.unique2)
   -&gt;  Seq Scan on tenk2 t2  (cost=0.00..458.00 rows=10000 width=244)
   -&gt;  Hash  (cost=232.35..232.35 rows=106 width=244)
         -&gt;  Bitmap Heap Scan on tenk1 t1  (cost=2.37..232.35 rows=106 width=244)
               Recheck Cond: (unique1 &lt; 100)
               -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..2.37 rows=106 width=0)
                     Index Cond: (unique1 &lt; 100)
</programlisting>

    This plan proposes to extract the 100 interesting rows of <classname>tenk1</classname>
    using that same old index scan, stash them into an in-memory hash table,
    and then do a sequential scan of <classname>tenk2</classname>, probing into the hash table
    for possible matches of <literal>t1.unique2 = t2.unique2</literal> for each <classname>tenk2</classname> row.
    The cost to read <classname>tenk1</classname> and set up the hash table is a start-up
    cost for the hash join, since there will be no output until we can
    start reading <classname>tenk2</classname>.  The total time estimate for the join also
    includes a hefty charge for the CPU time to probe the hash table
    10000 times.  Note, however, that we are <emphasis>not</emphasis> charging 10000 times 232.35;
    the hash table setup is only done once in this plan type.
   </para>

   <para>
    It is possible to check the accuracy of the planner's estimated costs
    by using <command>EXPLAIN ANALYZE</>.  This command actually executes the query,
    and then displays the true run time accumulated within each plan node
    along with the same estimated costs that a plain <command>EXPLAIN</command> shows.
    For example, we might get a result like this:

<screen>
EXPLAIN ANALYZE SELECT *
FROM tenk1 t1, tenk2 t2
WHERE t1.unique1 &lt; 100 AND t1.unique2 = t2.unique2;

                                                            QUERY PLAN
----------------------------------------------------------------------------------------------------------------------------------
 Nested Loop  (cost=2.37..553.11 rows=106 width=488) (actual time=1.392..12.700 rows=100 loops=1)
   -&gt;  Bitmap Heap Scan on tenk1 t1  (cost=2.37..232.35 rows=106 width=244) (actual time=0.878..2.367 rows=100 loops=1)
         Recheck Cond: (unique1 &lt; 100)
         -&gt;  Bitmap Index Scan on tenk1_unique1  (cost=0.00..2.37 rows=106 width=0) (actual time=0.546..0.546 rows=100 loops=1)
               Index Cond: (unique1 &lt; 100)
   -&gt;  Index Scan using tenk2_unique2 on tenk2 t2  (cost=0.00..3.01 rows=1 width=244) (actual time=0.067..0.078 rows=1 loops=100)
         Index Cond: (unique2 = t1.unique2)
 Total runtime: 14.452 ms
</screen>

    Note that the <quote>actual time</quote> values are in milliseconds of
    real time, whereas the <literal>cost</> estimates are expressed in
    arbitrary units; so they are unlikely to match up.
    The thing to pay attention to is whether the ratios of actual time and
    estimated costs are consistent.
   </para>

   <para>
    In some query plans, it is possible for a subplan node to be executed more
    than once.  For example, the inner index scan is executed once per outer
    row in the above nested-loop plan.  In such cases, the
    <literal>loops</> value reports the
    total number of executions of the node, and the actual time and rows
    values shown are averages per-execution.  This is done to make the numbers
    comparable with the way that the cost estimates are shown.  Multiply by
    the <literal>loops</> value to get the total time actually spent in
    the node.
   </para>

   <para>
    The <literal>Total runtime</literal> shown by <command>EXPLAIN
    ANALYZE</command> includes executor start-up and shut-down time, as well
    as time spent processing the result rows.  It does not include parsing,
    rewriting, or planning time.  For a <command>SELECT</> query, the total
    run time will normally be just a little larger than the total time
    reported for the top-level plan node.  For <command>INSERT</>,
    <command>UPDATE</>, and <command>DELETE</> commands, the total run time
    might be considerably larger, because it includes the time spent processing
    the result rows.  For these commands, the time for the top plan node is
    essentially the time spent locating the old rows and/or computing
    the new ones, but it doesn't include the time spent applying the changes.
    Time spent firing triggers, if any, is also outside the top plan node,
    and is shown separately for each trigger.
   </para>

   <para>
    It is worth noting that <command>EXPLAIN</> results should not be extrapolated
    to situations other than the one you are actually testing; for example,
    results on a toy-sized table cannot be assumed to apply to large tables.
    The planner's cost estimates are not linear and so it might choose
    a different plan for a larger or smaller table.  An extreme example
    is that on a table that only occupies one disk page, you'll nearly
    always get a sequential scan plan whether indexes are available or not.
    The planner realizes that it's going to take one disk page read to
    process the table in any case, so there's no value in expending additional
    page reads to look at an index.
   </para>
  </sect1>

 <sect1 id="planner-stats">
  <title>Statistics Used by the Planner</title>

  <indexterm zone="planner-stats">
   <primary>statistics</primary>
   <secondary>of the planner</secondary>
  </indexterm>

  <para>
   As we saw in the previous section, the query planner needs to estimate
   the number of rows retrieved by a query in order to make good choices
   of query plans.  This section provides a quick look at the statistics
   that the system uses for these estimates.
  </para>

  <para>
   One component of the statistics is the total number of entries in
   each table and index, as well as the number of disk blocks occupied
   by each table and index.  This information is kept in the table
   <link linkend="catalog-pg-class"><structname>pg_class</structname></link>,
   in the columns <structfield>reltuples</structfield> and
   <structfield>relpages</structfield>.  We can look at it with
   queries similar to this one:

<screen>
SELECT relname, relkind, reltuples, relpages
FROM pg_class
WHERE relname LIKE 'tenk1%';

       relname        | relkind | reltuples | relpages
----------------------+---------+-----------+----------
 tenk1                | r       |     10000 |      358
 tenk1_hundred        | i       |     10000 |       30
 tenk1_thous_tenthous | i       |     10000 |       30
 tenk1_unique1        | i       |     10000 |       30
 tenk1_unique2        | i       |     10000 |       30
(5 rows)
</screen>

   Here we can see that <structname>tenk1</structname> contains 10000
   rows, as do its indexes, but the indexes are (unsurprisingly) much
   smaller than the table.
  </para>

  <para>
   For efficiency reasons, <structfield>reltuples</structfield>
   and <structfield>relpages</structfield> are not updated on-the-fly,
   and so they usually contain somewhat out-of-date values.
   They are updated by <command>VACUUM</>, <command>ANALYZE</>, and a
   few DDL commands such as <command>CREATE INDEX</>.  A stand-alone
   <command>ANALYZE</>, that is one not part of <command>VACUUM</>,
   generates an approximate <structfield>reltuples</structfield> value
   since it does not read every row of the table.  The planner
   will scale the values it finds in <structname>pg_class</structname>
   to match the current physical table size, thus obtaining a closer
   approximation.
  </para>

  <indexterm>
   <primary>pg_statistic</primary>
  </indexterm>

  <para>
   Most queries retrieve only a fraction of the rows in a table, due
   to <literal>WHERE</> clauses that restrict the rows to be
   examined.  The planner thus needs to make an estimate of the
   <firstterm>selectivity</> of <literal>WHERE</> clauses, that is,
   the fraction of rows that match each condition in the
   <literal>WHERE</> clause.  The information used for this task is
   stored in the
   <link linkend="catalog-pg-statistic"><structname>pg_statistic</structname></link>
   system catalog.  Entries in <structname>pg_statistic</structname>
   are updated by the <command>ANALYZE</> and <command>VACUUM
   ANALYZE</> commands, and are always approximate even when freshly
   updated.
  </para>

  <indexterm>
   <primary>pg_stats</primary>
  </indexterm>

  <para>
   Rather than look at <structname>pg_statistic</structname> directly,
   it's better to look at its view
   <link linkend="view-pg-stats"><structname>pg_stats</structname></link>
   when examining the statistics manually.  <structname>pg_stats</structname>
   is designed to be more easily readable.  Furthermore,
   <structname>pg_stats</structname> is readable by all, whereas
   <structname>pg_statistic</structname> is only readable by a superuser.
   (This prevents unprivileged users from learning something about
   the contents of other people's tables from the statistics.  The
   <structname>pg_stats</structname> view is restricted to show only
   rows about tables that the current user can read.)
   For example, we might do:

<screen>
SELECT attname, inherited, n_distinct,
       array_to_string(most_common_vals, E'\n') as most_common_vals
FROM pg_stats
WHERE tablename = 'road';

 attname | inherited | n_distinct |          most_common_vals          
---------+-----------+------------+------------------------------------
 name    | f         |  -0.363388 | I- 580                        Ramp+
         |           |            | I- 880                        Ramp+
         |           |            | Sp Railroad                       +
         |           |            | I- 580                            +
         |           |            | I- 680                        Ramp
 name    | t         |  -0.284859 | I- 880                        Ramp+
         |           |            | I- 580                        Ramp+
         |           |            | I- 680                        Ramp+
         |           |            | I- 580                            +
         |           |            | State Hwy 13                  Ramp
(2 rows)
</screen>

   Note that two rows are displayed for the same column, one corresponding
   to the complete inheritance hierarchy starting at the
   <literal>road</literal> table (<literal>inherited</>=<literal>t</>),
   and another one including only the <literal>road</literal> table itself
   (<literal>inherited</>=<literal>f</>).
  </para>

  <para>
   The amount of information stored in <structname>pg_statistic</structname>
   by <command>ANALYZE</>, in particular the maximum number of entries in the
   <structfield>most_common_vals</> and <structfield>histogram_bounds</>
   arrays for each column, can be set on a
   column-by-column basis using the <command>ALTER TABLE SET STATISTICS</>
   command, or globally by setting the
   <xref linkend="guc-default-statistics-target"> configuration variable.
   The default limit is presently 100 entries.  Raising the limit
   might allow more accurate planner estimates to be made, particularly for
   columns with irregular data distributions, at the price of consuming
   more space in <structname>pg_statistic</structname> and slightly more
   time to compute the estimates.  Conversely, a lower limit might be
   sufficient for columns with simple data distributions.
  </para>

  <para>
   Further details about the planner's use of statistics can be found in
   <xref linkend="planner-stats-details">.
  </para>

 </sect1>

 <sect1 id="explicit-joins">
  <title>Controlling the Planner with Explicit <literal>JOIN</> Clauses</title>

  <indexterm zone="explicit-joins">
   <primary>join</primary>
   <secondary>controlling the order</secondary>
  </indexterm>

  <para>
   It is possible
   to control the query planner to some extent by using the explicit <literal>JOIN</>
   syntax.  To see why this matters, we first need some background.
  </para>

  <para>
   In a simple join query, such as:
<programlisting>
SELECT * FROM a, b, c WHERE a.id = b.id AND b.ref = c.id;
</programlisting>
   the planner is free to join the given tables in any order.  For
   example, it could generate a query plan that joins A to B, using
   the <literal>WHERE</> condition <literal>a.id = b.id</>, and then
   joins C to this joined table, using the other <literal>WHERE</>
   condition.  Or it could join B to C and then join A to that result.
   Or it could join A to C and then join them with B &mdash; but that
   would be inefficient, since the full Cartesian product of A and C
   would have to be formed, there being no applicable condition in the
   <literal>WHERE</> clause to allow optimization of the join.  (All
   joins in the <productname>PostgreSQL</productname> executor happen
   between two input tables, so it's necessary to build up the result
   in one or another of these fashions.)  The important point is that
   these different join possibilities give semantically equivalent
   results but might have hugely different execution costs.  Therefore,
   the planner will explore all of them to try to find the most
   efficient query plan.
  </para>

  <para>
   When a query only involves two or three tables, there aren't many join
   orders to worry about.  But the number of possible join orders grows
   exponentially as the number of tables expands.  Beyond ten or so input
   tables it's no longer practical to do an exhaustive search of all the
   possibilities, and even for six or seven tables planning might take an
   annoyingly long time.  When there are too many input tables, the
   <productname>PostgreSQL</productname> planner will switch from exhaustive
   search to a <firstterm>genetic</firstterm> probabilistic search
   through a limited number of possibilities.  (The switch-over threshold is
   set by the <xref linkend="guc-geqo-threshold"> run-time
   parameter.)
   The genetic search takes less time, but it won't
   necessarily find the best possible plan.
  </para>

  <para>
   When the query involves outer joins, the planner has less freedom
   than it does for plain (inner) joins. For example, consider:
<programlisting>
SELECT * FROM a LEFT JOIN (b JOIN c ON (b.ref = c.id)) ON (a.id = b.id);
</programlisting>
   Although this query's restrictions are superficially similar to the
   previous example, the semantics are different because a row must be
   emitted for each row of A that has no matching row in the join of B and C.
   Therefore the planner has no choice of join order here: it must join
   B to C and then join A to that result.  Accordingly, this query takes
   less time to plan than the previous query.  In other cases, the planner
   might be able to determine that more than one join order is safe.
   For example, given:
<programlisting>
SELECT * FROM a LEFT JOIN b ON (a.bid = b.id) LEFT JOIN c ON (a.cid = c.id);
</programlisting>
   it is valid to join A to either B or C first.  Currently, only
   <literal>FULL JOIN</> completely constrains the join order.  Most
   practical cases involving <literal>LEFT JOIN</> or <literal>RIGHT JOIN</>
   can be rearranged to some extent.
  </para>

  <para>
   Explicit inner join syntax (<literal>INNER JOIN</>, <literal>CROSS
   JOIN</>, or unadorned <literal>JOIN</>) is semantically the same as
   listing the input relations in <literal>FROM</>, so it does not
   constrain the join order.
  </para>

  <para>
   Even though most kinds of <literal>JOIN</> don't completely constrain
   the join order, it is possible to instruct the
   <productname>PostgreSQL</productname> query planner to treat all
   <literal>JOIN</> clauses as constraining the join order anyway.
   For example, these three queries are logically equivalent:
<programlisting>
SELECT * FROM a, b, c WHERE a.id = b.id AND b.ref = c.id;
SELECT * FROM a CROSS JOIN b CROSS JOIN c WHERE a.id = b.id AND b.ref = c.id;
SELECT * FROM a JOIN (b JOIN c ON (b.ref = c.id)) ON (a.id = b.id);
</programlisting>
   But if we tell the planner to honor the <literal>JOIN</> order,
   the second and third take less time to plan than the first.  This effect
   is not worth worrying about for only three tables, but it can be a
   lifesaver with many tables.
  </para>

  <para>
   To force the planner to follow the join order laid out by explicit
   <literal>JOIN</>s,
   set the <xref linkend="guc-join-collapse-limit"> run-time parameter to 1.
   (Other possible values are discussed below.)
  </para>

  <para>
   You do not need to constrain the join order completely in order to
   cut search time, because it's OK to use <literal>JOIN</> operators
   within items of a plain <literal>FROM</> list.  For example, consider:
<programlisting>
SELECT * FROM a CROSS JOIN b, c, d, e WHERE ...;
</programlisting>
   With <varname>join_collapse_limit</> = 1, this
   forces the planner to join A to B before joining them to other tables,
   but doesn't constrain its choices otherwise.  In this example, the
   number of possible join orders is reduced by a factor of 5.
  </para>

  <para>
   Constraining the planner's search in this way is a useful technique
   both for reducing planning time and for directing the planner to a
   good query plan.  If the planner chooses a bad join order by default,
   you can force it to choose a better order via <literal>JOIN</> syntax
   &mdash; assuming that you know of a better order, that is.  Experimentation
   is recommended.
  </para>

  <para>
   A closely related issue that affects planning time is collapsing of
   subqueries into their parent query.  For example, consider:
<programlisting>
SELECT *
FROM x, y,
    (SELECT * FROM a, b, c WHERE something) AS ss
WHERE somethingelse;
</programlisting>
   This situation might arise from use of a view that contains a join;
   the view's <literal>SELECT</> rule will be inserted in place of the view
   reference, yielding a query much like the above.  Normally, the planner
   will try to collapse the subquery into the parent, yielding:
<programlisting>
SELECT * FROM x, y, a, b, c WHERE something AND somethingelse;
</programlisting>
   This usually results in a better plan than planning the subquery
   separately.  (For example, the outer <literal>WHERE</> conditions might be such that
   joining X to A first eliminates many rows of A, thus avoiding the need to
   form the full logical output of the subquery.)  But at the same time,
   we have increased the planning time; here, we have a five-way join
   problem replacing two separate three-way join problems.  Because of the
   exponential growth of the number of possibilities, this makes a big
   difference.  The planner tries to avoid getting stuck in huge join search
   problems by not collapsing a subquery if more than <varname>from_collapse_limit</>
   <literal>FROM</> items would result in the parent
   query.  You can trade off planning time against quality of plan by
   adjusting this run-time parameter up or down.
  </para>

  <para>
   <xref linkend="guc-from-collapse-limit"> and <xref
   linkend="guc-join-collapse-limit">
   are similarly named because they do almost the same thing: one controls
   when the planner will <quote>flatten out</> subqueries, and the
   other controls when it will flatten out explicit joins.  Typically
   you would either set <varname>join_collapse_limit</> equal to
   <varname>from_collapse_limit</> (so that explicit joins and subqueries
   act similarly) or set <varname>join_collapse_limit</> to 1 (if you want
   to control join order with explicit joins).  But you might set them
   differently if you are trying to fine-tune the trade-off between planning
   time and run time.
  </para>
 </sect1>

 <sect1 id="populate">
  <title>Populating a Database</title>

  <para>
   One might need to insert a large amount of data when first populating
   a database. This section contains some suggestions on how to make
   this process as efficient as possible.
  </para>

  <sect2 id="disable-autocommit">
   <title>Disable Autocommit</title>

   <indexterm>
    <primary>autocommit</primary>
    <secondary>bulk-loading data</secondary>
   </indexterm>

   <para>
    When using multiple <command>INSERT</>s, turn off autocommit and just do
    one commit at the end.  (In plain
    SQL, this means issuing <command>BEGIN</command> at the start and
    <command>COMMIT</command> at the end.  Some client libraries might
    do this behind your back, in which case you need to make sure the
    library does it when you want it done.)  If you allow each
    insertion to be committed separately,
    <productname>PostgreSQL</productname> is doing a lot of work for
    each row that is added.  An additional benefit of doing all
    insertions in one transaction is that if the insertion of one row
    were to fail then the insertion of all rows inserted up to that
    point would be rolled back, so you won't be stuck with partially
    loaded data.
   </para>
  </sect2>

  <sect2 id="populate-copy-from">
   <title>Use <command>COPY</command></title>

   <para>
    Use <xref linkend="sql-copy"> to load
    all the rows in one command, instead of using a series of
    <command>INSERT</command> commands.  The <command>COPY</command>
    command is optimized for loading large numbers of rows; it is less
    flexible than <command>INSERT</command>, but incurs significantly
    less overhead for large data loads. Since <command>COPY</command>
    is a single command, there is no need to disable autocommit if you
    use this method to populate a table.
   </para>

   <para>
    If you cannot use <command>COPY</command>, it might help to use <xref
    linkend="sql-prepare"> to create a
    prepared <command>INSERT</command> statement, and then use
    <command>EXECUTE</command> as many times as required.  This avoids
    some of the overhead of repeatedly parsing and planning
    <command>INSERT</command>. Different interfaces provide this facility
    in different ways; look for <quote>prepared statements</> in the interface
    documentation.
   </para>

   <para>
    Note that loading a large number of rows using
    <command>COPY</command> is almost always faster than using
    <command>INSERT</command>, even if <command>PREPARE</> is used and
    multiple insertions are batched into a single transaction.
   </para>

   <para>
    <command>COPY</command> is fastest when used within the same
    transaction as an earlier <command>CREATE TABLE</command> or
    <command>TRUNCATE</command> command. In such cases no WAL
    needs to be written, because in case of an error, the files
    containing the newly loaded data will be removed anyway.
    However, this consideration only applies when
    <xref linkend="guc-wal-level"> is <literal>minimal</> as all commands
    must write WAL otherwise.
   </para>

  </sect2>

  <sect2 id="populate-rm-indexes">
   <title>Remove Indexes</title>

   <para>
    If you are loading a freshly created table, the fastest method is to
    create the table, bulk load the table's data using
    <command>COPY</command>, then create any indexes needed for the
    table.  Creating an index on pre-existing data is quicker than
    updating it incrementally as each row is loaded.
   </para>

   <para>
    If you are adding large amounts of data to an existing table,
    it might be a win to drop the indexes,
    load the table, and then recreate the indexes.  Of course, the
    database performance for other users might suffer
    during the time the indexes are missing.  One should also think
    twice before dropping a unique index, since the error checking
    afforded by the unique constraint will be lost while the index is
    missing.
   </para>
  </sect2>

  <sect2 id="populate-rm-fkeys">
   <title>Remove Foreign Key Constraints</title>

   <para>
    Just as with indexes, a foreign key constraint can be checked
    <quote>in bulk</> more efficiently than row-by-row.  So it might be
    useful to drop foreign key constraints, load data, and re-create
    the constraints.  Again, there is a trade-off between data load
    speed and loss of error checking while the constraint is missing.
   </para>

   <para>
    What's more, when you load data into a table with existing foreign key
    constraints, each new row requires an entry in the server's list of
    pending trigger events (since it is the firing of a trigger that checks
    the row's foreign key constraint).  Loading many millions of rows can
    cause the trigger event queue to overflow available memory, leading to
    intolerable swapping or even outright failure of the command.  Therefore
    it may be <emphasis>necessary</>, not just desirable, to drop and re-apply
    foreign keys when loading large amounts of data.  If temporarily removing
    the constraint isn't acceptable, the only other recourse may be to split
    up the load operation into smaller transactions.
   </para>
  </sect2>

  <sect2 id="populate-work-mem">
   <title>Increase <varname>maintenance_work_mem</varname></title>

   <para>
    Temporarily increasing the <xref linkend="guc-maintenance-work-mem">
    configuration variable when loading large amounts of data can
    lead to improved performance.  This will help to speed up <command>CREATE
    INDEX</> commands and <command>ALTER TABLE ADD FOREIGN KEY</> commands.
    It won't do much for <command>COPY</> itself, so this advice is
    only useful when you are using one or both of the above techniques.
   </para>
  </sect2>

  <sect2 id="populate-checkpoint-segments">
   <title>Increase <varname>checkpoint_segments</varname></title>

   <para>
    Temporarily increasing the <xref
    linkend="guc-checkpoint-segments"> configuration variable can also
    make large data loads faster.  This is because loading a large
    amount of data into <productname>PostgreSQL</productname> will
    cause checkpoints to occur more often than the normal checkpoint
    frequency (specified by the <varname>checkpoint_timeout</varname>
    configuration variable). Whenever a checkpoint occurs, all dirty
    pages must be flushed to disk. By increasing
    <varname>checkpoint_segments</varname> temporarily during bulk
    data loads, the number of checkpoints that are required can be
    reduced.
   </para>
  </sect2>

  <sect2 id="populate-pitr">
   <title>Disable WAL archival and streaming replication</title>

   <para>
    When loading large amounts of data into an installation that uses
    WAL archiving or streaming replication, it might be faster to take a
    new base backup after the load has completed than to process a large
    amount of incremental WAL data.  To prevent incremental WAL logging
    while loading, disable archiving and streaming replication, by setting
    <xref linkend="guc-wal-level"> to <literal>minimal</>,
    <xref linkend="guc-archive-mode"> to <literal>off</>, and
    <xref linkend="guc-max-wal-senders"> to zero.
    But note that changing these settings requires a server restart.
   </para>

   <para>
    Aside from avoiding the time for the archiver or WAL sender to
    process the WAL data,
    doing this will actually make certain commands faster, because they
    are designed not to write WAL at all if <varname>wal_level</varname>
    is <literal>minimal</>.  (They can guarantee crash safety more cheaply
    by doing an <function>fsync</> at the end than by writing WAL.)
    This applies to the following commands:
    <itemizedlist>
     <listitem>
      <para>
       <command>CREATE TABLE AS SELECT</command>
      </para>
     </listitem>
     <listitem>
      <para>
       <command>CREATE INDEX</command> (and variants such as
       <command>ALTER TABLE ADD PRIMARY KEY</command>)
      </para>
     </listitem>
     <listitem>
      <para>
       <command>ALTER TABLE SET TABLESPACE</command>
      </para>
     </listitem>
     <listitem>
      <para>
       <command>CLUSTER</command>
      </para>
     </listitem>
     <listitem>
      <para>
       <command>COPY FROM</command>, when the target table has been
       created or truncated earlier in the same transaction
      </para>
     </listitem>
    </itemizedlist>
   </para>
  </sect2>

  <sect2 id="populate-analyze">
   <title>Run <command>ANALYZE</command> Afterwards</title>

   <para>
    Whenever you have significantly altered the distribution of data
    within a table, running <xref linkend="sql-analyze"> is strongly recommended. This
    includes bulk loading large amounts of data into the table.  Running
    <command>ANALYZE</command> (or <command>VACUUM ANALYZE</command>)
    ensures that the planner has up-to-date statistics about the
    table.  With no statistics or obsolete statistics, the planner might
    make poor decisions during query planning, leading to poor
    performance on any tables with inaccurate or nonexistent
    statistics.  Note that if the autovacuum daemon is enabled, it might
    run <command>ANALYZE</command> automatically; see
    <xref linkend="vacuum-for-statistics">
    and <xref linkend="autovacuum"> for more information.
   </para>
  </sect2>

  <sect2 id="populate-pg-dump">
   <title>Some Notes About <application>pg_dump</></title>

   <para>
    Dump scripts generated by <application>pg_dump</> automatically apply
    several, but not all, of the above guidelines.  To reload a
    <application>pg_dump</> dump as quickly as possible, you need to
    do a few extra things manually.  (Note that these points apply while
    <emphasis>restoring</> a dump, not while <emphasis>creating</> it.
    The same points apply whether loading a text dump with
    <application>psql</> or using <application>pg_restore</> to load
    from a <application>pg_dump</> archive file.)
   </para>

   <para>
    By default, <application>pg_dump</> uses <command>COPY</>, and when
    it is generating a complete schema-and-data dump, it is careful to
    load data before creating indexes and foreign keys.  So in this case
    several guidelines are handled automatically.  What is left
    for you to do is to:
    <itemizedlist>
     <listitem>
      <para>
       Set appropriate (i.e., larger than normal) values for
       <varname>maintenance_work_mem</varname> and
       <varname>checkpoint_segments</varname>.
      </para>
     </listitem>
     <listitem>
      <para>
       If using WAL archiving or streaming replication, consider disabling
       them during the restore. To do that, set <varname>archive_mode</>
       to <literal>off</>,
       <varname>wal_level</varname> to <literal>minimal</>, and
       <varname>max_wal_senders</> to zero before loading the dump.
       Afterwards, set them back to the right values and take a fresh
       base backup.
      </para>
     </listitem>
     <listitem>
      <para>
       Consider whether the whole dump should be restored as a single
       transaction.  To do that, pass the <option>-1</> or
       <option>--single-transaction</> command-line option to
       <application>psql</> or <application>pg_restore</>. When using this
       mode, even the smallest of errors will rollback the entire restore,
       possibly discarding many hours of processing.  Depending on how
       interrelated the data is, that might seem preferable to manual cleanup,
       or not.  <command>COPY</> commands will run fastest if you use a single
       transaction and have WAL archiving turned off.
      </para>
     </listitem>
     <listitem>
      <para>
       If multiple CPUs are available in the database server, consider using
       <application>pg_restore</>'s <option>--jobs</> option.  This
       allows concurrent data loading and index creation.
      </para>
     </listitem>
     <listitem>
      <para>
       Run <command>ANALYZE</> afterwards.
      </para>
     </listitem>
    </itemizedlist>
   </para>

   <para>
    A data-only dump will still use <command>COPY</>, but it does not
    drop or recreate indexes, and it does not normally touch foreign
    keys.

     <footnote>
      <para>
       You can get the effect of disabling foreign keys by using
       the <option>--disable-triggers</> option &mdash; but realize that
       that eliminates, rather than just postpones, foreign key
       validation, and so it is possible to insert bad data if you use it.
      </para>
     </footnote>

    So when loading a data-only dump, it is up to you to drop and recreate
    indexes and foreign keys if you wish to use those techniques.
    It's still useful to increase <varname>checkpoint_segments</varname>
    while loading the data, but don't bother increasing
    <varname>maintenance_work_mem</varname>; rather, you'd do that while
    manually recreating indexes and foreign keys afterwards.
    And don't forget to <command>ANALYZE</> when you're done; see
    <xref linkend="vacuum-for-statistics">
    and <xref linkend="autovacuum"> for more information.
   </para>
  </sect2>
  </sect1>

  <sect1 id="non-durability">
   <title>Non-Durable Settings</title>

   <indexterm zone="non-durability">
    <primary>non-durable</primary>
   </indexterm>

   <para>
    Durability is a database feature that guarantees the recording of
    committed transactions even if the server crashes or loses
    power.  However, durability adds significant database overhead,
    so if your site does not require such a guarantee,
    <productname>PostgreSQL</productname> can be configured to run
    much faster.  The following are configuration changes you can make
    to improve performance in such cases;  they do not invalidate
    commit guarantees related to database crashes, only abrupt operating
    system stoppage, except as mentioned below:

    <itemizedlist>
     <listitem>
      <para>
       Place the database cluster's data directory in a memory-backed
       file system (i.e. <acronym>RAM</> disk).  This eliminates all
       database disk I/O, but limits data storage to the amount of
       available memory (and perhaps swap).
      </para>
     </listitem>

     <listitem>
      <para>
       Turn off <xref linkend="guc-fsync">;  there is no need to flush
       data to disk.
      </para>
     </listitem>

     <listitem>
      <para>
       Turn off <xref linkend="guc-full-page-writes">;  there is no need
       to guard against partial page writes.
      </para>
     </listitem>

     <listitem>
      <para>
       Increase <xref linkend="guc-checkpoint-segments"> and <xref
       linkend="guc-checkpoint-timeout"> ; this reduces the frequency
       of checkpoints, but increases the storage requirements of
       <filename>/pg_xlog</>.
      </para>
     </listitem>

     <listitem>
      <para>
       Turn off <xref linkend="guc-synchronous-commit">;  there might be no
       need to write the <acronym>WAL</acronym> to disk on every
       commit.  This does affect database crash transaction durability.
      </para>
     </listitem>
    </itemizedlist>
   </para>
  </sect1>

 </chapter>
