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$PostgreSQL: pgsql/doc/src/sgml/perform.sgml,v 1.50 2005/02/03 07:12:37 neilc Exp $
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 <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 manipulated 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 is given.  Choosing the right
    plan to match the query structure and the properties of the data
    is absolutely critical for good performance.  You can use the
    <xref linkend="sql-explain" endterm="sql-explain-title"> command
    to see what query plan the system 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 numbers that are currently quoted by <command>EXPLAIN</command> are:

    <itemizedlist>
     <listitem>
      <para>
       Estimated start-up cost (Time expended before 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 were to be retrieved, which they may not
       be: a query with a <literal>LIMIT</> clause will stop short of paying the total cost,
       for example.)
      </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 units of disk page fetches.  (CPU effort
    estimates are converted into disk-page units using some
    fairly arbitrary fudge factors. If you want to experiment with these
    factors, see the list of run-time configuration parameters in
    <xref linkend="runtime-config-query-constants">.)
   </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/optimizer cares about.
    In particular, the cost does not consider the time spent transmitting
    result rows to the frontend, which could be a pretty dominant
    factor in the true 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>
    Rows output is a little tricky because it is <emphasis>not</emphasis> the
    number of rows
    processed/scanned by the query, it is usually less, reflecting the
    estimated selectivity of any <literal>WHERE</>-clause conditions that are being
    applied at this node.  Ideally the top-level rows estimate will
    approximate the number of rows actually returned, updated, or deleted
    by the query.
   </para>

   <para>
    Here are some examples (using the regression test database after a
    <command>VACUUM ANALYZE</>, and 7.3 development sources):

<programlisting>
EXPLAIN SELECT * FROM tenk1;

                         QUERY PLAN
-------------------------------------------------------------
 Seq Scan on tenk1  (cost=0.00..333.00 rows=10000 width=148)
</programlisting>
   </para>

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

<programlisting>
SELECT * FROM pg_class WHERE relname = 'tenk1';
</programlisting>

    you will find out that <classname>tenk1</classname> has 233 disk
    pages and 10000 rows.  So the cost is estimated at 233 page
    reads, defined as costing 1.0 apiece, plus 10000 * <xref
    linkend="guc-cpu-tuple-cost"> which is
    currently 0.01 (try <command>SHOW cpu_tuple_cost</command>).
   </para>

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

<programlisting>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 1000;

                         QUERY PLAN
------------------------------------------------------------
 Seq Scan on tenk1  (cost=0.00..358.00 rows=1033 width=148)
   Filter: (unique1 &lt; 1000)
</programlisting>

    The estimate of output rows has gone down 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 to reflect the extra CPU
    time spent checking the <literal>WHERE</> condition.
   </para>

   <para>
    The actual number of rows this query would select is 1000, but the
    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>
    Modify the query to restrict the condition even more:

<programlisting>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 50;

                                   QUERY PLAN
-------------------------------------------------------------------------------
 Index Scan using tenk1_unique1 on tenk1  (cost=0.00..179.33 rows=49 width=148)
   Index Cond: (unique1 &lt; 50)
</programlisting>

    and you will see that if we make the <literal>WHERE</> condition selective
    enough, the planner will
    eventually decide that an index scan is cheaper than a sequential scan.
    This plan will only have to visit 50 rows because of the index,
    so it wins despite the fact that each individual fetch is more expensive
    than reading a whole disk page sequentially.
   </para>

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

<programlisting>
EXPLAIN SELECT * FROM tenk1 WHERE unique1 &lt; 50 AND stringu1 = 'xxx';

                                  QUERY PLAN
-------------------------------------------------------------------------------
 Index Scan using tenk1_unique1 on tenk1  (cost=0.00..179.45 rows=1 width=148)
   Index Cond: (unique1 &lt; 50)
   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
    a little bit to reflect this extra checking.
   </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; 50 AND t1.unique2 = t2.unique2;

                               QUERY PLAN
----------------------------------------------------------------------------
 Nested Loop  (cost=0.00..327.02 rows=49 width=296)
   -&gt;  Index Scan using tenk1_unique1 on tenk1 t1
                                      (cost=0.00..179.33 rows=49 width=148)
         Index Cond: (unique1 &lt; 50)
   -&gt;  Index Scan using tenk2_unique2 on tenk2 t2
                                      (cost=0.00..3.01 rows=1 width=148)
         Index Cond: ("outer".unique2 = t2.unique2)
</programlisting>
   </para>

   <para>
    In this nested-loop join, the outer scan is the same index scan we had
    in the example before last, and so its cost and row count are the same
    because we are applying the <literal>WHERE</> clause <literal>unique1 &lt; 50</literal> at that node.
    The <literal>t1.unique2 = t2.unique2</literal> clause is not relevant yet, so it doesn't
    affect row count of the outer scan.  For the inner 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>t2.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 (49 * 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 general, because
    in general 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 winner, using the enable/disable
    flags for each plan type.  (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; 50 AND t1.unique2 = t2.unique2;

                               QUERY PLAN
--------------------------------------------------------------------------
 Hash Join  (cost=179.45..563.06 rows=49 width=296)
   Hash Cond: ("outer".unique2 = "inner".unique2)
   -&gt;  Seq Scan on tenk2 t2  (cost=0.00..333.00 rows=10000 width=148)
   -&gt;  Hash  (cost=179.33..179.33 rows=49 width=148)
         -&gt;  Index Scan using tenk1_unique1 on tenk1 t1
                                    (cost=0.00..179.33 rows=49 width=148)
               Index Cond: (unique1 &lt; 50)
</programlisting>

    This plan proposes to extract the 50 interesting rows of <classname>tenk1</classname>
    using ye same olde 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> at each <classname>tenk2</classname> row.
    The cost to read <classname>tenk1</classname> and set up the hash table is entirely start-up
    cost for the hash join, since we won't get any rows out 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 179.33;
    the hash table setup is only done once in this plan type.
   </para>

   <para>
    It is possible to check on 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; 50 AND t1.unique2 = t2.unique2;

                                   QUERY PLAN
-------------------------------------------------------------------------------
 Nested Loop  (cost=0.00..327.02 rows=49 width=296)
                                 (actual time=1.181..29.822 rows=50 loops=1)
   -&gt;  Index Scan using tenk1_unique1 on tenk1 t1
                  (cost=0.00..179.33 rows=49 width=148)
                                 (actual time=0.630..8.917 rows=50 loops=1)
         Index Cond: (unique1 &lt; 50)
   -&gt;  Index Scan using tenk2_unique2 on tenk2 t2
                  (cost=0.00..3.01 rows=1 width=148)
                                 (actual time=0.295..0.324 rows=1 loops=50)
         Index Cond: ("outer".unique2 = t2.unique2)
 Total runtime: 31.604 ms
</screen>

    Note that the <quote>actual time</quote> values are in milliseconds of
    real time, whereas the <quote>cost</quote> estimates are expressed in
    arbitrary units of disk fetches; so they are unlikely to match up.
    The thing to pay attention to is the ratios.
   </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
    <quote>loops</quote> 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 <quote>loops</quote> 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 may be
    considerably larger, because it includes the time spent processing the
    result rows.  In these commands, the time for the top plan node
    essentially is the time spent computing the new rows and/or locating
    the old ones, but it doesn't include the time spent making the changes.
   </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 can't be assumed to apply to large tables.
    The planner's cost estimates are not linear and so it may well 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 |      233
 tenk1_hundred | i       |     10000 |       30
 tenk1_unique1 | i       |     10000 |       30
 tenk1_unique2 | i       |     10000 |       30
(4 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 having <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 <structname>pg_stats</structname>
   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, n_distinct, most_common_vals FROM pg_stats WHERE tablename = 'road';

 attname | n_distinct |                                                                                                                                                                                  most_common_vals                                                                                                                                                                                   
---------+------------+-------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------
 name    |  -0.467008 | {"I- 580                        Ramp","I- 880                        Ramp","Sp Railroad                       ","I- 580                            ","I- 680                        Ramp","I- 80                         Ramp","14th                          St  ","5th                           St  ","Mission                       Blvd","I- 880                            "}
 thepath |         20 | {"[(-122.089,37.71),(-122.0886,37.711)]"}
(2 rows)
</screen>
  </para>

  <para>
   <structname>pg_stats</structname> is described in detail in
   <xref linkend="view-pg-stats">.
  </para>

  <para>
   The amount of information stored in <structname>pg_statistic</structname>,
   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 10 entries.  Raising the limit
   may 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 may be
   appropriate for columns with simple data distributions.
  </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, 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 may 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 may 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 much 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.
  </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 need to
   constrain the join order.  But it is possible to instruct the
   <productname>PostgreSQL</productname> query planner to treat
   explicit inner <literal>JOIN</>s 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 <literal>JOIN</> order for inner joins,
   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</> subselects, and the
   other controls when it will flatten out explicit inner joins.  Typically
   you would either set <varname>join_collapse_limit</> equal to
   <varname>from_collapse_limit</> (so that explicit joins and subselects
   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 may 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>
    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 may
    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" endterm="sql-copy-title"> 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 may help to use <xref
    linkend="sql-prepare" endterm="sql-prepare-title"> 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>.
   </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>
  </sect2>

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

   <para>
    If you are loading a freshly created table, the fastest way 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 augmenting an existing table, you can drop the index,
    load the table, and then recreate the index. Of course, the
    database performance for other users may be adversely affected
    during the time that the index is missing.  One should also think
    twice before dropping unique indexes, since the error checking
    afforded by the unique constraint will be lost while the index is
    missing.
   </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 is because when a B-tree index
    is created from scratch, the existing content of the table needs
    to be sorted. Allowing the merge sort to use more memory
    means that fewer merge passes will be required.  A larger setting for
    <varname>maintenance_work_mem</varname> may also speed up validation
    of foreign-key constraints.
   </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> can
    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-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"
    endterm="sql-analyze-title"> 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 may
    make poor decisions during query planning, leading to poor
    performance on any tables with inaccurate or nonexistent
    statistics.
   </para>
  </sect2>
  </sect1>

 </chapter>

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