tree-ssa.texi

理解和实践操作系统的一本好书

     4	  x_10 = @dots{}
     5	  if (x_1 > 7)
     6	    y_2 = 0
     7	  else
     8	    x_11 = @dots{}
     9	    y_3 = x_1 + x_7
     10	  endif
     11	  x_5 = x_1 + 1
     12	  goto L0;
     13	endif
@end smallexample

We want to replace all the uses of @code{x_1} with the new definitions
of @code{x_10} and @code{x_11}.  Note that the only uses that should
be replaced are those at lines @code{5}, @code{9} and @code{11}.
Also, the use of @code{x_7} at line @code{9} should @emph{not} be
replaced (this is why we cannot just mark symbol @code{x} for
renaming)@.

Additionally, we may need to insert a PHI node at line @code{11}
because that is a merge point for @code{x_10} and @code{x_11}.  So the
use of @code{x_1} at line @code{11} will be replaced with the new PHI
node.  The insertion of PHI nodes is optional.  They are not strictly
necessary to preserve the SSA form, and depending on what the caller
inserted, they may not even be useful for the optimizers@.

Updating the SSA form is a two step process.  First, the pass has to
identify which names need to be updated and/or which symbols need to
be renamed into SSA form for the first time.  When new names are
introduced to replace existing names in the program, the mapping
between the old and the new names are registered by calling
@code{register_new_name_mapping} (note that if your pass creates new
code by duplicating basic blocks, the call to @code{tree_duplicate_bb}
will set up the necessary mappings automatically).  On the other hand,
if your pass exposes a new symbol that should be put in SSA form for
the first time, the new symbol should be registered with
@code{mark_sym_for_renaming}.

After the replacement mappings have been registered and new symbols
marked for renaming, a call to @code{update_ssa} makes the registered
changes.  This can be done with an explicit call or by creating
@code{TODO} flags in the @code{tree_opt_pass} structure for your pass.
There are several @code{TODO} flags that control the behavior of
@code{update_ssa}:

@itemize @bullet
@item @code{TODO_update_ssa}.  Update the SSA form inserting PHI nodes
      for newly exposed symbols and virtual names marked for updating.
      When updating real names, only insert PHI nodes for a real name
      @code{O_j} in blocks reached by all the new and old definitions for
      @code{O_j}.  If the iterated dominance frontier for @code{O_j}
      is not pruned, we may end up inserting PHI nodes in blocks that
      have one or more edges with no incoming definition for
      @code{O_j}.  This would lead to uninitialized warnings for
      @code{O_j}'s symbol@.

@item @code{TODO_update_ssa_no_phi}.  Update the SSA form without
      inserting any new PHI nodes at all.  This is used by passes that
      have either inserted all the PHI nodes themselves or passes that
      need only to patch use-def and def-def chains for virtuals
      (e.g., DCE)@.


@item @code{TODO_update_ssa_full_phi}.  Insert PHI nodes everywhere
      they are needed.  No pruning of the IDF is done.  This is used
      by passes that need the PHI nodes for @code{O_j} even if it
      means that some arguments will come from the default definition
      of @code{O_j}'s symbol (e.g., @code{pass_linear_transform})@.

      WARNING: If you need to use this flag, chances are that your
      pass may be doing something wrong.  Inserting PHI nodes for an
      old name where not all edges carry a new replacement may lead to
      silent codegen errors or spurious uninitialized warnings@.

@item @code{TODO_update_ssa_only_virtuals}.  Passes that update the
      SSA form on their own may want to delegate the updating of
      virtual names to the generic updater.  Since FUD chains are
      easier to maintain, this simplifies the work they need to do.
      NOTE: If this flag is used, any OLD->NEW mappings for real names
      are explicitly destroyed and only the symbols marked for
      renaming are processed@.
@end itemize

@subsection Preserving the virtual SSA form
@cindex preserving virtual SSA form

The virtual SSA form is harder to preserve than the non-virtual SSA form
mainly because the set of virtual operands for a statement may change at
what some would consider unexpected times.  In general, statement
modifications should be bracketed between calls to
@code{push_stmt_changes} and @code{pop_stmt_changes}.  For example,

@smallexample
    munge_stmt (tree stmt)
    @{
       push_stmt_changes (&stmt);
       @dots{} rewrite STMT @dots{}
       pop_stmt_changes (&stmt);
    @}
@end smallexample

The call to @code{push_stmt_changes} saves the current state of the
statement operands and the call to @code{pop_stmt_changes} compares
the saved state with the current one and does the appropriate symbol
marking for the SSA renamer.

It is possible to modify several statements at a time, provided that
@code{push_stmt_changes} and @code{pop_stmt_changes} are called in
LIFO order, as when processing a stack of statements.

Additionally, if the pass discovers that it did not need to make
changes to the statement after calling @code{push_stmt_changes}, it
can simply discard the topmost change buffer by calling
@code{discard_stmt_changes}.  This will avoid the expensive operand
re-scan operation and the buffer comparison that determines if symbols
need to be marked for renaming.

@subsection Examining @code{SSA_NAME} nodes
@cindex examining SSA_NAMEs

The following macros can be used to examine @code{SSA_NAME} nodes

@defmac SSA_NAME_DEF_STMT (@var{var})
Returns the statement @var{s} that creates the @code{SSA_NAME}
@var{var}.  If @var{s} is an empty statement (i.e., @code{IS_EMPTY_STMT
(@var{s})} returns @code{true}), it means that the first reference to
this variable is a USE or a VUSE@.
@end defmac

@defmac SSA_NAME_VERSION (@var{var})
Returns the version number of the @code{SSA_NAME} object @var{var}.
@end defmac


@subsection Walking use-def chains

@deftypefn {Tree SSA function} void walk_use_def_chains (@var{var}, @var{fn}, @var{data})

Walks use-def chains starting at the @code{SSA_NAME} node @var{var}.
Calls function @var{fn} at each reaching definition found.  Function
@var{FN} takes three arguments: @var{var}, its defining statement
(@var{def_stmt}) and a generic pointer to whatever state information
that @var{fn} may want to maintain (@var{data}).  Function @var{fn} is
able to stop the walk by returning @code{true}, otherwise in order to
continue the walk, @var{fn} should return @code{false}.

Note, that if @var{def_stmt} is a @code{PHI} node, the semantics are
slightly different.  For each argument @var{arg} of the PHI node, this
function will:

@enumerate
@item	Walk the use-def chains for @var{arg}.
@item	Call @code{FN (@var{arg}, @var{phi}, @var{data})}.
@end enumerate

Note how the first argument to @var{fn} is no longer the original
variable @var{var}, but the PHI argument currently being examined.
If @var{fn} wants to get at @var{var}, it should call
@code{PHI_RESULT} (@var{phi}).
@end deftypefn

@subsection Walking the dominator tree

@deftypefn {Tree SSA function} void walk_dominator_tree (@var{walk_data}, @var{bb})

This function walks the dominator tree for the current CFG calling a
set of callback functions defined in @var{struct dom_walk_data} in
@file{domwalk.h}.  The call back functions you need to define give you
hooks to execute custom code at various points during traversal:

@enumerate
@item Once to initialize any local data needed while processing
      @var{bb} and its children.  This local data is pushed into an
      internal stack which is automatically pushed and popped as the
      walker traverses the dominator tree.

@item Once before traversing all the statements in the @var{bb}.

@item Once for every statement inside @var{bb}.

@item Once after traversing all the statements and before recursing
      into @var{bb}'s dominator children.

@item It then recurses into all the dominator children of @var{bb}.

@item After recursing into all the dominator children of @var{bb} it
      can, optionally, traverse every statement in @var{bb} again
      (i.e., repeating steps 2 and 3).

@item Once after walking the statements in @var{bb} and @var{bb}'s
      dominator children.  At this stage, the block local data stack
      is popped.
@end enumerate
@end deftypefn

@node Alias analysis
@section Alias analysis
@cindex alias
@cindex flow-sensitive alias analysis
@cindex flow-insensitive alias analysis

Alias analysis proceeds in 4 main phases:

@enumerate
@item   Structural alias analysis.

This phase walks the types for structure variables, and determines which
of the fields can overlap using offset and size of each field.  For each
field, a ``subvariable'' called a ``Structure field tag'' (SFT)@ is
created, which represents that field as a separate variable.  All
accesses that could possibly overlap with a given field will have
virtual operands for the SFT of that field.

@smallexample
struct foo
@{
  int a;
  int b;
@}
struct foo temp;
int bar (void)
@{
  int tmp1, tmp2, tmp3;
  SFT.0_2 = VDEF <SFT.0_1>
  temp.a = 5;
  SFT.1_4 = VDEF <SFT.1_3>
  temp.b = 6;
  
  VUSE <SFT.1_4>
  tmp1_5 = temp.b;
  VUSE <SFT.0_2>
  tmp2_6 = temp.a;

  tmp3_7 = tmp1_5 + tmp2_6;
  return tmp3_7;
@}
@end smallexample

If you copy the symbol tag for a variable for some reason, you probably
also want to copy the subvariables for that variable.

@item	Points-to and escape analysis.

This phase walks the use-def chains in the SSA web looking for
three things:

	@itemize @bullet
	@item	Assignments of the form @code{P_i = &VAR}
	@item	Assignments of the form P_i = malloc()
	@item	Pointers and ADDR_EXPR that escape the current function.
	@end itemize

The concept of `escaping' is the same one used in the Java world.
When a pointer or an ADDR_EXPR escapes, it means that it has been
exposed outside of the current function.  So, assignment to
global variables, function arguments and returning a pointer are
all escape sites.

This is where we are currently limited.  Since not everything is
renamed into SSA, we lose track of escape properties when a
pointer is stashed inside a field in a structure, for instance.
In those cases, we are assuming that the pointer does escape.

We use escape analysis to determine whether a variable is
call-clobbered.  Simply put, if an ADDR_EXPR escapes, then the
variable is call-clobbered.  If a pointer P_i escapes, then all
the variables pointed-to by P_i (and its memory tag) also escape.

@item	Compute flow-sensitive aliases

We have two classes of memory tags.  Memory tags associated with
the pointed-to data type of the pointers in the program.  These
tags are called ``symbol memory tag'' (SMT)@.  The other class are
those associated with SSA_NAMEs, called ``name memory tag'' (NMT)@.
The basic idea is that when adding operands for an INDIRECT_REF
*P_i, we will first check whether P_i has a name tag, if it does
we use it, because that will have more precise aliasing
information.  Otherwise, we use the standard symbol tag.

In this phase, we go through all the pointers we found in
points-to analysis and create alias sets for the name memory tags
associated with each pointer P_i.  If P_i escapes, we mark
call-clobbered the variables it points to and its tag.


@item	Compute flow-insensitive aliases

This pass will compare the alias set of every symbol memory tag and
every addressable variable found in the program.  Given a symbol
memory tag SMT and an addressable variable V@.  If the alias sets
of SMT and V conflict (as computed by may_alias_p), then V is
marked as an alias tag and added to the alias set of SMT@.

Every language that wishes to perform language-specific alias analysis
should define a function that computes, given a @code{tree}
node, an alias set for the node.  Nodes in different alias sets are not
allowed to alias.  For an example, see the C front-end function
@code{c_get_alias_set}.
@end enumerate

For instance, consider the following function:

@smallexample
foo (int i)
@{
  int *p, *q, a, b;

  if (i > 10)
    p = &a;
  else
    q = &b;

  *p = 3;
  *q = 5;
  a = b + 2;
  return *p;
@}
@end smallexample

After aliasing analysis has finished, the symbol memory tag for
pointer @code{p} will have two aliases, namely variables @code{a} and
@code{b}.
Every time pointer @code{p} is dereferenced, we want to mark the
operation as a potential reference to @code{a} and @code{b}.

@smallexample
foo (int i)
@{
  int *p, a, b;

  if (i_2 > 10)
    p_4 = &a;
  else
    p_6 = &b;
  # p_1 = PHI <p_4(1), p_6(2)>;

  # a_7 = VDEF <a_3>;
  # b_8 = VDEF <b_5>;
  *p_1 = 3;

  # a_9 = VDEF <a_7>
  # VUSE <b_8>
  a_9 = b_8 + 2;

  # VUSE <a_9>;
  # VUSE <b_8>;
  return *p_1;
@}
@end smallexample

In certain cases, the list of may aliases for a pointer may grow
too large.  This may cause an explosion in the number of virtual
operands inserted in the code.  Resulting in increased memory
consumption and compilation time.

When the number of virtual operands needed to represent aliased
loads and stores grows too large (configurable with @option{--param
max-aliased-vops}), alias sets are grouped to avoid severe
compile-time slow downs and memory consumption.  The alias
grouping heuristic proceeds as follows:

@enumerate
@item Sort the list of pointers in decreasing number of contributed
virtual operands.

@item Take the first pointer from the list and reverse the role
of the memory tag and its aliases.  Usually, whenever an
aliased variable Vi is found to alias with a memory tag
T, we add Vi to the may-aliases set for T@.  Meaning that
after alias analysis, we will have:

@smallexample
may-aliases(T) = @{ V1, V2, V3, @dots{}, Vn @}
@end smallexample

This means that every statement that references T, will get
@code{n} virtual operands for each of the Vi tags.  But, when
alias grouping is enabled, we make T an alias tag and add it
to the alias set of all the Vi variables:

@smallexample
may-aliases(V1) = @{ T @}
may-aliases(V2) = @{ T @}
@dots{}
may-aliases(Vn) = @{ T @}
@end smallexample

This has two effects: (a) statements referencing T will only get
a single virtual operand, and, (b) all the variables Vi will now
appear to alias each other.  So, we lose alias precision to
improve compile time.  But, in theory, a program with such a high
level of aliasing should not be very optimizable in the first
place.

@item Since variables may be in the alias set of more than one
memory tag, the grouping done in step (2) needs to be extended
to all the memory tags that have a non-empty intersection with
the may-aliases set of tag T@.  For instance, if we originally
had these may-aliases sets:

@smallexample
may-aliases(T) = @{ V1, V2, V3 @}
may-aliases(R) = @{ V2, V4 @}
@end smallexample

In step (2) we would have reverted the aliases for T as:

@smallexample
may-aliases(V1) = @{ T @}
may-aliases(V2) = @{ T @}
may-aliases(V3) = @{ T @}
@end smallexample

But note that now V2 is no longer aliased with R@.  We could
add R to may-aliases(V2), but we are in the process of
grouping aliases to reduce virtual operands so what we do is
add V4 to the grouping to obtain:

@smallexample
may-aliases(V1) = @{ T @}
may-aliases(V2) = @{ T @}
may-aliases(V3) = @{ T @}
may-aliases(V4) = @{ T @}
@end smallexample

@item If the total number of virtual operands due to aliasing is
still above the threshold set by max-alias-vops, go back to (2).
@end enumerate