tree-ssa.texi   [plain text]

@c Copyright (c) 2004, 2005 Free Software Foundation, Inc.
@c Free Software Foundation, Inc.
@c This is part of the GCC manual.
@c For copying conditions, see the file gcc.texi.

@c ---------------------------------------------------------------------
@c Tree SSA
@c ---------------------------------------------------------------------

@node Tree SSA
@chapter Analysis and Optimization of GIMPLE Trees
@cindex Tree SSA
@cindex Optimization infrastructure for GIMPLE

GCC uses three main intermediate languages to represent the program
during compilation: GENERIC, GIMPLE and RTL@.  GENERIC is a
language-independent representation generated by each front end.  It
is used to serve as an interface between the parser and optimizer.
GENERIC is a common representation that is able to represent programs
written in all the languages supported by GCC@.

GIMPLE and RTL are used to optimize the program.  GIMPLE is used for
target and language independent optimizations (e.g., inlining,
constant propagation, tail call elimination, redundancy elimination,
etc).  Much like GENERIC, GIMPLE is a language independent, tree based
representation.  However, it differs from GENERIC in that the GIMPLE
grammar is more restrictive: expressions contain no more than 3
operands (except function calls), it has no control flow structures
and expressions with side-effects are only allowed on the right hand
side of assignments.  See the chapter describing GENERIC and GIMPLE
for more details.

This chapter describes the data structures and functions used in the
GIMPLE optimizers (also known as ``tree optimizers'' or ``middle
end'').  In particular, it focuses on all the macros, data structures,
functions and programming constructs needed to implement optimization
passes for GIMPLE@.

* GENERIC::		A high-level language-independent representation.
* GIMPLE::              A lower-level factored tree representation.
* Annotations::		Attributes for statements and variables.
* Statement Operands::	Variables referenced by GIMPLE statements.
* SSA::			Static Single Assignment representation.
* Alias analysis::	Representing aliased loads and stores.
@end menu

@section GENERIC
@cindex GENERIC

The purpose of GENERIC is simply to provide a language-independent way of
representing an entire function in trees.  To this end, it was necessary to
add a few new tree codes to the back end, but most everything was already
there.  If you can express it with the codes in @code{gcc/tree.def}, it's

Early on, there was a great deal of debate about how to think about
statements in a tree IL@.  In GENERIC, a statement is defined as any
expression whose value, if any, is ignored.  A statement will always
have @code{TREE_SIDE_EFFECTS} set (or it will be discarded), but a
non-statement expression may also have side effects.  A
@code{CALL_EXPR}, for instance.

It would be possible for some local optimizations to work on the
GENERIC form of a function; indeed, the adapted tree inliner works
fine on GENERIC, but the current compiler performs inlining after
lowering to GIMPLE (a restricted form described in the next section).
Indeed, currently the frontends perform this lowering before handing
off to @code{tree_rest_of_compilation}, but this seems inelegant.

If necessary, a front end can use some language-dependent tree codes
in its GENERIC representation, so long as it provides a hook for
converting them to GIMPLE and doesn't expect them to work with any
(hypothetical) optimizers that run before the conversion to GIMPLE@.
The intermediate representation used while parsing C and C++ looks
very little like GENERIC, but the C and C++ gimplifier hooks are
perfectly happy to take it as input and spit out GIMPLE@.

@node GIMPLE
@section GIMPLE
@cindex GIMPLE

GIMPLE is a simplified subset of GENERIC for use in optimization.  The
particular subset chosen (and the name) was heavily influenced by the
SIMPLE IL used by the McCAT compiler project at McGill University,
though we have made some different choices.  For one thing, SIMPLE
doesn't support @code{goto}; a production compiler can't afford that
kind of restriction.

GIMPLE retains much of the structure of the parse trees: lexical
scopes are represented as containers, rather than markers.  However,
expressions are broken down into a 3-address form, using temporary
variables to hold intermediate values.  Also, control structures are
lowered to gotos.

In GIMPLE no container node is ever used for its value; if a
@code{COND_EXPR} or @code{BIND_EXPR} has a value, it is stored into a
temporary within the controlled blocks, and that temporary is used in
place of the container.

The compiler pass which lowers GENERIC to GIMPLE is referred to as the
@samp{gimplifier}.  The gimplifier works recursively, replacing complex
statements with sequences of simple statements.

@c Currently, the only way to
@c tell whether or not an expression is in GIMPLE form is by recursively
@c examining it; in the future there will probably be a flag to help avoid
@c redundant work.  FIXME FIXME

* Interfaces::
* Temporaries::
* GIMPLE Expressions::
* Statements::
* GIMPLE Example::
* Rough GIMPLE Grammar::
@end menu

@node Interfaces
@subsection Interfaces
@cindex gimplification

The tree representation of a function is stored in
@code{DECL_SAVED_TREE}.  It is lowered to GIMPLE by a call to

If a front end wants to include language-specific tree codes in the tree
representation which it provides to the back end, it must provide a
definition of @code{LANG_HOOKS_GIMPLIFY_EXPR} which knows how to
convert the front end trees to GIMPLE@.  Usually such a hook will involve
much of the same code for expanding front end trees to RTL@.  This function
can return fully lowered GIMPLE, or it can return GENERIC trees and let the
main gimplifier lower them the rest of the way; this is often simpler.
GIMPLE that is not fully lowered is known as ``high GIMPLE'' and
consists of the IL before the pass @code{pass_lower_cf}.  High GIMPLE
still contains lexical scopes and nested expressions, while low GIMPLE
exposes all of the implicit jumps for control expressions like

The C and C++ front ends currently convert directly from front end
trees to GIMPLE, and hand that off to the back end rather than first
converting to GENERIC@.  Their gimplifier hooks know about all the
@code{_STMT} nodes and how to convert them to GENERIC forms.  There
was some work done on a genericization pass which would run first, but
the existence of @code{STMT_EXPR} meant that in order to convert all
of the C statements into GENERIC equivalents would involve walking the
entire tree anyway, so it was simpler to lower all the way.  This
might change in the future if someone writes an optimization pass
which would work better with higher-level trees, but currently the
optimizers all expect GIMPLE@.

A front end which wants to use the tree optimizers (and already has
some sort of whole-function tree representation) only needs to provide
a definition of @code{LANG_HOOKS_GIMPLIFY_EXPR}, call
@code{gimplify_function_tree} to lower to GIMPLE, and then hand off to
@code{tree_rest_of_compilation} to compile and output the function.

You can tell the compiler to dump a C-like representation of the GIMPLE
form with the flag @option{-fdump-tree-gimple}.

@node Temporaries
@subsection Temporaries
@cindex Temporaries

When gimplification encounters a subexpression which is too complex, it
creates a new temporary variable to hold the value of the subexpression,
and adds a new statement to initialize it before the current statement.
These special temporaries are known as @samp{expression temporaries}, and are
allocated using @code{get_formal_tmp_var}.  The compiler tries to
always evaluate identical expressions into the same temporary, to simplify
elimination of redundant calculations.

We can only use expression temporaries when we know that it will not be
reevaluated before its value is used, and that it will not be otherwise
modified@footnote{These restrictions are derived from those in Morgan 4.8.}.
Other temporaries can be allocated using
@code{get_initialized_tmp_var} or @code{create_tmp_var}.

Currently, an expression like @code{a = b + 5} is not reduced any
further.  We tried converting it to something like
  T1 = b + 5;
  a = T1;
@end smallexample
but this bloated the representation for minimal benefit.  However, a
variable which must live in memory cannot appear in an expression; its
value is explicitly loaded into a temporary first.  Similarly, storing
the value of an expression to a memory variable goes through a

@node GIMPLE Expressions
@subsection Expressions
@cindex GIMPLE Expressions

In general, expressions in GIMPLE consist of an operation and the
appropriate number of simple operands; these operands must either be a
GIMPLE rvalue (@code{is_gimple_val}), i.e.@: a constant or a register
variable.  More complex operands are factored out into temporaries, so
  a = b + c + d
@end smallexample
  T1 = b + c;
  a = T1 + d;
@end smallexample

The same rule holds for arguments to a @code{CALL_EXPR}.

The target of an assignment is usually a variable, but can also be an
@code{INDIRECT_REF} or a compound lvalue as described below.

* Compound Expressions::
* Compound Lvalues::
* Conditional Expressions::
* Logical Operators::
@end menu

@node Compound Expressions
@subsubsection Compound Expressions
@cindex Compound Expressions

The left-hand side of a C comma expression is simply moved into a separate

@node Compound Lvalues
@subsubsection Compound Lvalues
@cindex Compound Lvalues

Currently compound lvalues involving array and structure field references
are not broken down; an expression like @code{a.b[2] = 42} is not reduced
any further (though complex array subscripts are).  This restriction is a
workaround for limitations in later optimizers; if we were to convert this

  T1 = &a.b;
  T1[2] = 42;
@end smallexample

alias analysis would not remember that the reference to @code{T1[2]} came
by way of @code{a.b}, so it would think that the assignment could alias
another member of @code{a}; this broke @code{struct-alias-1.c}.  Future
optimizer improvements may make this limitation unnecessary.

@node Conditional Expressions
@subsubsection Conditional Expressions
@cindex Conditional Expressions

A C @code{?:} expression is converted into an @code{if} statement with
each branch assigning to the same temporary.  So,

  a = b ? c : d;
@end smallexample
  if (b)
    T1 = c;
    T1 = d;
  a = T1;
@end smallexample

Tree level if-conversion pass re-introduces @code{?:} expression, if appropriate.
It is used to vectorize loops with conditions using vector conditional operations.

Note that in GIMPLE, @code{if} statements are also represented using
@code{COND_EXPR}, as described below.

@node Logical Operators
@subsubsection Logical Operators
@cindex Logical Operators

Except when they appear in the condition operand of a @code{COND_EXPR},
logical `and' and `or' operators are simplified as follows:
@code{a = b && c} becomes

  T1 = (bool)b;
  if (T1)
    T1 = (bool)c;
  a = T1;
@end smallexample

Note that @code{T1} in this example cannot be an expression temporary,
because it has two different assignments.

@node Statements
@subsection Statements
@cindex Statements

Most statements will be assignment statements, represented by
@code{MODIFY_EXPR}.  A @code{CALL_EXPR} whose value is ignored can
also be a statement.  No other C expressions can appear at statement level;
a reference to a volatile object is converted into a @code{MODIFY_EXPR}.
In GIMPLE form, type of @code{MODIFY_EXPR} is not meaningful.  Instead, use type
of LHS or RHS@.

There are also several varieties of complex statements.

* Blocks::
* Statement Sequences::
* Empty Statements::
* Loops::
* Selection Statements::
* Jumps::
* Cleanups::
* GIMPLE Exception Handling::
@end menu

@node Blocks
@subsubsection Blocks
@cindex Blocks

Block scopes and the variables they declare in GENERIC and GIMPLE are
expressed using the @code{BIND_EXPR} code, which in previous versions of
GCC was primarily used for the C statement-expression extension.

Variables in a block are collected into @code{BIND_EXPR_VARS} in
declaration order.  Any runtime initialization is moved out of
@code{DECL_INITIAL} and into a statement in the controlled block.  When
gimplifying from C or C++, this initialization replaces the

Variable-length arrays (VLAs) complicate this process, as their size often
refers to variables initialized earlier in the block.  To handle this, we
currently split the block at that point, and move the VLA into a new, inner
@code{BIND_EXPR}.  This strategy may change in the future.

@code{DECL_SAVED_TREE} for a GIMPLE function will always be a
@code{BIND_EXPR} which contains declarations for the temporary variables
used in the function.

A C++ program will usually contain more @code{BIND_EXPR}s than there are
syntactic blocks in the source code, since several C++ constructs have
implicit scopes associated with them.  On the other hand, although the C++
front end uses pseudo-scopes to handle cleanups for objects with
destructors, these don't translate into the GIMPLE form; multiple
declarations at the same level use the same @code{BIND_EXPR}.

@node Statement Sequences
@subsubsection Statement Sequences
@cindex Statement Sequences

Multiple statements at the same nesting level are collected into a
@code{STATEMENT_LIST}.  Statement lists are modified and traversed
using the interface in @samp{tree-iterator.h}.

@node Empty Statements
@subsubsection Empty Statements
@cindex Empty Statements

Whenever possible, statements with no effect are discarded.  But if they
are nested within another construct which cannot be discarded for some
reason, they are instead replaced with an empty statement, generated by
@code{build_empty_stmt}.  Initially, all empty statements were shared,
after the pattern of the Java front end, but this caused a lot of trouble in

An empty statement is represented as @code{(void)0}.

@node Loops
@subsubsection Loops
@cindex Loops

At one time loops were expressed in GIMPLE using @code{LOOP_EXPR}, but
now they are lowered to explicit gotos.

@node Selection Statements
@subsubsection Selection Statements
@cindex Selection Statements

A simple selection statement, such as the C @code{if} statement, is
expressed in GIMPLE using a void @code{COND_EXPR}.  If only one branch is
used, the other is filled with an empty statement.

Normally, the condition expression is reduced to a simple comparison.  If
it is a shortcut (@code{&&} or @code{||}) expression, however, we try to
break up the @code{if} into multiple @code{if}s so that the implied shortcut
is taken directly, much like the transformation done by @code{do_jump} in
the RTL expander.

A @code{SWITCH_EXPR} in GIMPLE contains the condition and a
@code{TREE_VEC} of @code{CASE_LABEL_EXPR}s describing the case values
and corresponding @code{LABEL_DECL}s to jump to.  The body of the
@code{switch} is moved after the @code{SWITCH_EXPR}.

@node Jumps
@subsubsection Jumps
@cindex Jumps

Other jumps are expressed by either @code{GOTO_EXPR} or @code{RETURN_EXPR}.

The operand of a @code{GOTO_EXPR} must be either a label or a variable
containing the address to jump to.

The operand of a @code{RETURN_EXPR} is either @code{NULL_TREE},
@code{RESULT_DECL}, or a @code{MODIFY_EXPR} which sets the return value.  It
would be nice to move the @code{MODIFY_EXPR} into a separate statement, but the
special return semantics in @code{expand_return} make that difficult.  It may
still happen in the future, perhaps by moving most of that logic into

@node Cleanups
@subsubsection Cleanups
@cindex Cleanups

Destructors for local C++ objects and similar dynamic cleanups are
represented in GIMPLE by a @code{TRY_FINALLY_EXPR}.
@code{TRY_FINALLY_EXPR} has two operands, both of which are a sequence
of statements to execute.  The first sequence is executed.  When it
completes the second sequence is executed.

The first sequence may complete in the following ways:


@item Execute the last statement in the sequence and fall off the

@item Execute a goto statement (@code{GOTO_EXPR}) to an ordinary
label outside the sequence.

@item Execute a return statement (@code{RETURN_EXPR}).

@item Throw an exception.  This is currently not explicitly represented in

@end enumerate

The second sequence is not executed if the first sequence completes by
calling @code{setjmp} or @code{exit} or any other function that does
not return.  The second sequence is also not executed if the first
sequence completes via a non-local goto or a computed goto (in general
the compiler does not know whether such a goto statement exits the
first sequence or not, so we assume that it doesn't).

After the second sequence is executed, if it completes normally by
falling off the end, execution continues wherever the first sequence
would have continued, by falling off the end, or doing a goto, etc.

@code{TRY_FINALLY_EXPR} complicates the flow graph, since the cleanup
needs to appear on every edge out of the controlled block; this
reduces the freedom to move code across these edges.  Therefore, the
EH lowering pass which runs before most of the optimization passes
eliminates these expressions by explicitly adding the cleanup to each
edge.  Rethrowing the exception is represented using @code{RESX_EXPR}.

@node GIMPLE Exception Handling
@subsubsection Exception Handling
@cindex GIMPLE Exception Handling

Other exception handling constructs are represented using
@code{TRY_CATCH_EXPR}.  @code{TRY_CATCH_EXPR} has two operands.  The
first operand is a sequence of statements to execute.  If executing
these statements does not throw an exception, then the second operand
is ignored.  Otherwise, if an exception is thrown, then the second
operand of the @code{TRY_CATCH_EXPR} is checked.  The second operand
may have the following forms:


@item A sequence of statements to execute.  When an exception occurs,
these statements are executed, and then the exception is rethrown.

@item A sequence of @code{CATCH_EXPR} expressions.  Each @code{CATCH_EXPR}
has a list of applicable exception types and handler code.  If the
thrown exception matches one of the caught types, the associated
handler code is executed.  If the handler code falls off the bottom,
execution continues after the original @code{TRY_CATCH_EXPR}.

@item An @code{EH_FILTER_EXPR} expression.  This has a list of
permitted exception types, and code to handle a match failure.  If the
thrown exception does not match one of the allowed types, the
associated match failure code is executed.  If the thrown exception
does match, it continues unwinding the stack looking for the next

@end enumerate

Currently throwing an exception is not directly represented in GIMPLE,
since it is implemented by calling a function.  At some point in the future
we will want to add some way to express that the call will throw an
exception of a known type.

Just before running the optimizers, the compiler lowers the high-level
EH constructs above into a set of @samp{goto}s, magic labels, and EH
regions.  Continuing to unwind at the end of a cleanup is represented
with a @code{RESX_EXPR}.

@node GIMPLE Example
@subsection GIMPLE Example
@cindex GIMPLE Example

struct A @{ A(); ~A(); @};

int i;
int g();
void f()
  A a;
  int j = (--i, i ? 0 : 1);

  for (int x = 42; x > 0; --x)
      i += g()*4 + 32;
@end smallexample


void f()
  int i.0;
  int T.1;
  int iftmp.2;
  int T.3;
  int T.4;
  int T.5;
  int T.6;

    struct A a;
    int j;

    __comp_ctor (&a);
        i.0 = i;
        T.1 = i.0 - 1;
        i = T.1;
        i.0 = i;
        if (i.0 == 0)
          iftmp.2 = 1;
          iftmp.2 = 0;
        j = iftmp.2;
          int x;

          x = 42;
          goto test;

          T.3 = g ();
          T.4 = T.3 * 4;
          i.0 = i;
          T.5 = T.4 + i.0;
          T.6 = T.5 + 32;
          i = T.6;
          x = x - 1;

          if (x > 0)
            goto loop;
            goto break_;
        __comp_dtor (&a);
@end smallexample

@node Rough GIMPLE Grammar
@subsection Rough GIMPLE Grammar
@cindex Rough GIMPLE Grammar

   function     : FUNCTION_DECL
                        DECL_SAVED_TREE -> compound-stmt

   compound-stmt: STATEMENT_LIST
                        members -> stmt

   stmt         : block
                | if-stmt
                | switch-stmt
                | goto-stmt
                | return-stmt
                | resx-stmt
                | label-stmt
                | try-stmt
                | modify-stmt
                | call-stmt

   block        : BIND_EXPR
                        BIND_EXPR_VARS -> chain of DECLs
                        BIND_EXPR_BLOCK -> BLOCK
                        BIND_EXPR_BODY -> compound-stmt

   if-stmt      : COND_EXPR
                        op0 -> condition
                        op1 -> compound-stmt
                        op2 -> compound-stmt

   switch-stmt  : SWITCH_EXPR
                        op0 -> val
                        op1 -> NULL
                        op2 -> TREE_VEC of CASE_LABEL_EXPRs
                            The CASE_LABEL_EXPRs are sorted by CASE_LOW,
                            and default is last.

   goto-stmt    : GOTO_EXPR
                        op0 -> LABEL_DECL | val

   return-stmt  : RETURN_EXPR
                        op0 -> return-value

   return-value : NULL
                | RESULT_DECL
                | MODIFY_EXPR
                        op0 -> RESULT_DECL
                        op1 -> lhs

   resx-stmt    : RESX_EXPR

   label-stmt   : LABEL_EXPR
                        op0 -> LABEL_DECL

   try-stmt     : TRY_CATCH_EXPR
                        op0 -> compound-stmt
                        op1 -> handler
                | TRY_FINALLY_EXPR
                        op0 -> compound-stmt
                        op1 -> compound-stmt

   handler      : catch-seq
                | EH_FILTER_EXPR
                | compound-stmt

   catch-seq    : STATEMENT_LIST
                        members -> CATCH_EXPR

   modify-stmt  : MODIFY_EXPR
                        op0 -> lhs
                        op1 -> rhs

   call-stmt    : CALL_EXPR
                        op0 -> val | OBJ_TYPE_REF
                        op1 -> call-arg-list

   call-arg-list: TREE_LIST
                        members -> lhs | CONST

   addr-expr-arg: ID
                | compref

   addressable  : addr-expr-arg
                | indirectref

   with-size-arg: addressable
                | call-stmt

   indirectref  : INDIRECT_REF
                        op0 -> val

   lhs          : addressable
                | bitfieldref
                | WITH_SIZE_EXPR
                        op0 -> with-size-arg
                        op1 -> val

   min-lval     : ID
                | indirectref

   bitfieldref  : BIT_FIELD_REF
                        op0 -> inner-compref
                        op1 -> CONST
                        op2 -> var

   compref      : inner-compref
                | TARGET_MEM_REF
                        op0 -> ID
                        op1 -> val
                        op2 -> val
                        op3 -> CONST
                        op4 -> CONST
                | REALPART_EXPR
                        op0 -> inner-compref
                | IMAGPART_EXPR
                        op0 -> inner-compref

   inner-compref: min-lval
                | COMPONENT_REF
                        op0 -> inner-compref
                        op1 -> FIELD_DECL
                        op2 -> val
                | ARRAY_REF
                        op0 -> inner-compref
                        op1 -> val
                        op2 -> val
                        op3 -> val
                | ARRAY_RANGE_REF
                        op0 -> inner-compref
                        op1 -> val
                        op2 -> val
                        op3 -> val
                | VIEW_CONVERT_EXPR
                        op0 -> inner-compref

   condition    : val
                | RELOP
                        op0 -> val
                        op1 -> val

   val          : ID
                | CONST

   rhs          : lhs
                | CONST
                | call-stmt
                | ADDR_EXPR
                        op0 -> addr-expr-arg
                | UNOP
                        op0 -> val
                | BINOP
                        op0 -> val
                        op1 -> val
                | RELOP
                        op0 -> val
                        op1 -> val
			op0 -> condition
			op1 -> val
			op2 -> val
@end smallexample

@node Annotations
@section Annotations
@cindex annotations

The optimizers need to associate attributes with statements and
variables during the optimization process.  For instance, we need to
know what basic block a statement belongs to or whether a variable
has aliases.  All these attributes are stored in data structures
called annotations which are then linked to the field @code{ann} in
@code{struct tree_common}.

Presently, we define annotations for statements (@code{stmt_ann_t}),
variables (@code{var_ann_t}) and SSA names (@code{ssa_name_ann_t}).
Annotations are defined and documented in @file{tree-flow.h}.

@node Statement Operands
@section Statement Operands
@cindex operands
@cindex virtual operands
@cindex real operands
@findex update_stmt

Almost every GIMPLE statement will contain a reference to a variable
or memory location.  Since statements come in different shapes and
sizes, their operands are going to be located at various spots inside
the statement's tree.  To facilitate access to the statement's
operands, they are organized into lists associated inside each
statement's annotation.  Each element in an operand list is a pointer
to a @code{VAR_DECL}, @code{PARM_DECL} or @code{SSA_NAME} tree node.
This provides a very convenient way of examining and replacing

Data flow analysis and optimization is done on all tree nodes
representing variables.  Any node for which @code{SSA_VAR_P} returns
nonzero is considered when scanning statement operands.  However, not
all @code{SSA_VAR_P} variables are processed in the same way.  For the
purposes of optimization, we need to distinguish between references to
local scalar variables and references to globals, statics, structures,
arrays, aliased variables, etc.  The reason is simple, the compiler
can gather complete data flow information for a local scalar.  On the
other hand, a global variable may be modified by a function call, it
may not be possible to keep track of all the elements of an array or
the fields of a structure, etc.

The operand scanner gathers two kinds of operands: @dfn{real} and
@dfn{virtual}.  An operand for which @code{is_gimple_reg} returns true
is considered real, otherwise it is a virtual operand.  We also
distinguish between uses and definitions.  An operand is used if its
value is loaded by the statement (e.g., the operand at the RHS of an
assignment).  If the statement assigns a new value to the operand, the
operand is considered a definition (e.g., the operand at the LHS of
an assignment).

Virtual and real operands also have very different data flow
properties.  Real operands are unambiguous references to the
full object that they represent.  For instance, given

  int a, b;
  a = b
@end smallexample

Since @code{a} and @code{b} are non-aliased locals, the statement
@code{a = b} will have one real definition and one real use because
variable @code{b} is completely modified with the contents of
variable @code{a}.  Real definition are also known as @dfn{killing
definitions}.  Similarly, the use of @code{a} reads all its bits.

In contrast, virtual operands are used with variables that can have
a partial or ambiguous reference.  This includes structures, arrays,
globals, and aliased variables.  In these cases, we have two types of
definitions.  For globals, structures, and arrays, we can determine from
a statement whether a variable of these types has a killing definition.
If the variable does, then the statement is marked as having a
@dfn{must definition} of that variable.  However, if a statement is only
defining a part of the variable (i.e.@: a field in a structure), or if we
know that a statement might define the variable but we cannot say for sure,
then we mark that statement as having a @dfn{may definition}.  For
instance, given

  int a, b, *p;

  if (...)
    p = &a;
    p = &b;
  *p = 5;
  return *p;
@end smallexample

The assignment @code{*p = 5} may be a definition of @code{a} or
@code{b}.  If we cannot determine statically where @code{p} is
pointing to at the time of the store operation, we create virtual
definitions to mark that statement as a potential definition site for
@code{a} and @code{b}.  Memory loads are similarly marked with virtual
use operands.  Virtual operands are shown in tree dumps right before
the statement that contains them.  To request a tree dump with virtual
operands, use the @option{-vops} option to @option{-fdump-tree}:

  int a, b, *p;

  if (...)
    p = &a;
    p = &b;
  # a = V_MAY_DEF <a>
  # b = V_MAY_DEF <b>
  *p = 5;

  # VUSE <a>
  # VUSE <b>
  return *p;
@end smallexample

Notice that @code{V_MAY_DEF} operands have two copies of the referenced
variable.  This indicates that this is not a killing definition of
that variable.  In this case we refer to it as a @dfn{may definition}
or @dfn{aliased store}.  The presence of the second copy of the
variable in the @code{V_MAY_DEF} operand will become important when the
function is converted into SSA form.  This will be used to link all
the non-killing definitions to prevent optimizations from making
incorrect assumptions about them.

Operands are updated as soon as the statement is finished via a call
to @code{update_stmt}.  If statement elements are changed via
@code{SET_USE} or @code{SET_DEF}, then no further action is required
(i.e., those macros take care of updating the statement).  If changes
are made by manipulating the statement's tree directly, then a call
must be made to @code{update_stmt} when complete.  Calling one of the
@code{bsi_insert} routines or @code{bsi_replace} performs an implicit
call to @code{update_stmt}.

@subsection Operand Iterators And Access Routines
@cindex Operand Iterators 
@cindex Operand Access Routines

Operands are collected by @file{tree-ssa-operands.c}.  They are stored
inside each statement's annotation and can be accessed through either the
operand iterators or an access routine.

The following access routines are available for examining operands:

@item @code{SINGLE_SSA_@{USE,DEF,TREE@}_OPERAND}: These accessors will return 
NULL unless there is exactly one operand matching the specified flags.  If 
there is exactly one operand, the operand is returned as either a @code{tree}, 
@code{def_operand_p}, or @code{use_operand_p}.

tree t = SINGLE_SSA_TREE_OPERAND (stmt, flags);
def_operand_p d = SINGLE_SSA_DEF_OPERAND (stmt, SSA_OP_ALL_DEFS);
@end smallexample

@item @code{ZERO_SSA_OPERANDS}: This macro returns true if there are no 
operands matching the specified flags.

@end smallexample

@item @code{NUM_SSA_OPERANDS}: This macro Returns the number of operands 
matching 'flags'.  This actually executes a loop to perform the count, so 
only use this if it is really needed.

int count = NUM_SSA_OPERANDS (stmt, flags)
@end smallexample
@end enumerate

If you wish to iterate over some or all operands, use the
@code{FOR_EACH_SSA_@{USE,DEF,TREE@}_OPERAND} iterator.  For example, to print
all the operands for a statement:

print_ops (tree stmt)
  tree var;

    print_generic_expr (stderr, var, TDF_SLIM);
@end smallexample

How to choose the appropriate iterator:

@item Determine whether you are need to see the operand pointers, or just the
    trees, and choose the appropriate macro:

Need            Macro:
----            -------
use_operand_p   FOR_EACH_SSA_USE_OPERAND
def_operand_p   FOR_EACH_SSA_DEF_OPERAND
@end smallexample

@item You need to declare a variable of the type you are interested
    in, and an ssa_op_iter structure which serves as the loop
    controlling variable.

@item Determine which operands you wish to use, and specify the flags of
    those you are interested in.  They are documented in

#define SSA_OP_USE              0x01    /* @r{Real USE operands.}  */
#define SSA_OP_DEF              0x02    /* @r{Real DEF operands.}  */
#define SSA_OP_VUSE             0x04    /* @r{VUSE operands.}  */
#define SSA_OP_VMAYUSE          0x08    /* @r{USE portion of V_MAY_DEFS.}  */
#define SSA_OP_VMAYDEF          0x10    /* @r{DEF portion of V_MAY_DEFS.}  */
#define SSA_OP_VMUSTDEF         0x20    /* @r{V_MUST_DEF definitions.}  */

/* @r{These are commonly grouped operand flags.}  */
@end smallexample
@end enumerate

So if you want to look at the use pointers for all the @code{USE} and
@code{VUSE} operands, you would do something like:

  use_operand_p use_p;
  ssa_op_iter iter;

  FOR_EACH_SSA_USE_OPERAND (use_p, stmt, iter, (SSA_OP_USE | SSA_OP_VUSE))
      process_use_ptr (use_p);
@end smallexample

The @code{TREE} macro is basically the same as the @code{USE} and
@code{DEF} macros, only with the use or def dereferenced via
@code{USE_FROM_PTR (use_p)} and @code{DEF_FROM_PTR (def_p)}.  Since we
aren't using operand pointers, use and defs flags can be mixed.

  tree var;
  ssa_op_iter iter;

       print_generic_expr (stderr, var, TDF_SLIM);
@end smallexample

@code{V_MAY_DEF}s are broken into two flags, one for the
@code{DEF} portion (@code{SSA_OP_VMAYDEF}) and one for the USE portion
(@code{SSA_OP_VMAYUSE}).  If all you want to look at are the
@code{V_MAY_DEF}s together, there is a fourth iterator macro for this,
which returns both a def_operand_p and a use_operand_p for each
@code{V_MAY_DEF} in the statement.  Note that you don't need any flags for
this one.

  use_operand_p use_p;
  def_operand_p def_p;
  ssa_op_iter iter;

  FOR_EACH_SSA_MAYDEF_OPERAND (def_p, use_p, stmt, iter)
@end smallexample

@code{V_MUST_DEF}s are broken into two flags, one for the
@code{DEF} portion (@code{SSA_OP_VMUSTDEF}) and one for the kill portion
(@code{SSA_OP_VMUSTKILL}).  If all you want to look at are the
@code{V_MUST_DEF}s together, there is a fourth iterator macro for this,
which returns both a def_operand_p and a use_operand_p for each
@code{V_MUST_DEF} in the statement.  Note that you don't need any flags for
this one.

  use_operand_p kill_p;
  def_operand_p def_p;
  ssa_op_iter iter;

  FOR_EACH_SSA_MUSTDEF_OPERAND (def_p, kill_p, stmt, iter)
@end smallexample

There are many examples in the code as well, as well as the
documentation in @file{tree-ssa-operands.h}.

There are also a couple of variants on the stmt iterators regarding PHI

@code{FOR_EACH_PHI_ARG} Works exactly like 
@code{FOR_EACH_SSA_USE_OPERAND}, except it works over @code{PHI} arguments 
instead of statement operands.

/* Look at every virtual PHI use.  */
FOR_EACH_PHI_ARG (use_p, phi_stmt, iter, SSA_OP_VIRTUAL_USES)

/* Look at every real PHI use.  */
FOR_EACH_PHI_ARG (use_p, phi_stmt, iter, SSA_OP_USES)

/* Look at every every PHI use.  */
FOR_EACH_PHI_ARG (use_p, phi_stmt, iter, SSA_OP_ALL_USES)
@end smallexample

@code{FOR_EACH_PHI_OR_STMT_@{USE,DEF@}} works exactly like 
@code{FOR_EACH_SSA_@{USE,DEF@}_OPERAND}, except it will function on
either a statement or a @code{PHI} node.  These should be used when it is
appropriate but they are not quite as efficient as the individual 
@code{FOR_EACH_PHI} and @code{FOR_EACH_SSA} routines.

FOR_EACH_PHI_OR_STMT_USE (use_operand_p, stmt, iter, flags)

FOR_EACH_PHI_OR_STMT_DEF (def_operand_p, phi, iter, flags)
@end smallexample

@subsection Immediate Uses
@cindex Immediate Uses

Immediate use information is now always available.  Using the immediate use 
iterators, you may examine every use of any @code{SSA_NAME}. For instance,
to change each use of @code{ssa_var} to @code{ssa_var2} and call fold_stmt on
each stmt after that is done:

  use_operand_p imm_use_p;
  imm_use_iterator iterator;
  tree ssa_var, stmt;

  FOR_EACH_IMM_USE_STMT (stmt, iterator, ssa_var)
      FOR_EACH_IMM_USE_ON_STMT (imm_use_p, iterator)
        SET_USE (imm_use_p, ssa_var_2);
      fold_stmt (stmt);
@end smallexample

There are 2 iterators which can be used. @code{FOR_EACH_IMM_USE_FAST} is
used when the immediate uses are not changed, i.e., you are looking at the
uses, but not setting them.  

If they do get changed, then care must be taken that things are not changed 
under the iterators, so use the @code{FOR_EACH_IMM_USE_STMT} and 
@code{FOR_EACH_IMM_USE_ON_STMT} iterators.  They attempt to preserve the 
sanity of the use list by moving all the uses for a statement into 
a controlled position, and then iterating over those uses.  Then the 
optimization can manipulate the stmt when all the uses have been
processed.  This is a little slower than the FAST version since it adds a 
placeholder element and must sort through the list a bit for each statement.  
This placeholder element must be also be removed if the loop is 
terminated early.  The macro @code{BREAK_FROM_IMM_USE_SAFE} is provided 
to do this :

  FOR_EACH_IMM_USE_STMT (stmt, iterator, ssa_var)
      if (stmt == last_stmt)
        BREAK_FROM_SAFE_IMM_USE (iter);

      FOR_EACH_IMM_USE_ON_STMT (imm_use_p, iterator)
        SET_USE (imm_use_p, ssa_var_2);
      fold_stmt (stmt);
@end smallexample

There are checks in @code{verify_ssa} which verify that the immediate use list
is up to date, as well as checking that an optimization didn't break from the 
loop without using this macro.  It is safe to simply 'break'; from a 
@code{FOR_EACH_IMM_USE_FAST} traverse.

Some useful functions and macros:
@item  @code{has_zero_uses (ssa_var)} : Returns true if there are no uses of
@item   @code{has_single_use (ssa_var)} : Returns true if there is only a 
single use of @code{ssa_var}.
@item   @code{single_imm_use (ssa_var, use_operand_p *ptr, tree *stmt)} :
Returns true if there is only a single use of @code{ssa_var}, and also returns
the use pointer and statement it occurs in in the second and third parameters.
@item   @code{num_imm_uses (ssa_var)} : Returns the number of immediate uses of
@code{ssa_var}. It is better not to use this if possible since it simply
utilizes a loop to count the uses.
@item  @code{PHI_ARG_INDEX_FROM_USE (use_p)} : Given a use within a @code{PHI}
node, return the index number for the use.  An assert is triggered if the use
isn't located in a @code{PHI} node.
@item  @code{USE_STMT (use_p)} : Return the statement a use occurs in.
@end enumerate

Note that uses are not put into an immediate use list until their statement is
actually inserted into the instruction stream via a @code{bsi_*} routine.  

It is also still possible to utilize lazy updating of statements, but this 
should be used only when absolutely required.  Both alias analysis and the 
dominator optimizations currently do this.  

When lazy updating is being used, the immediate use information is out of date 
and cannot be used reliably.  Lazy updating is achieved by simply marking
statements modified via calls to @code{mark_stmt_modified} instead of 
@code{update_stmt}.  When lazy updating is no longer required, all the 
modified statements must have @code{update_stmt} called in order to bring them 
up to date.  This must be done before the optimization is finished, or 
@code{verify_ssa} will trigger an abort.

This is done with a simple loop over the instruction stream:
  block_stmt_iterator bsi;
  basic_block bb;
  FOR_EACH_BB (bb)
      for (bsi = bsi_start (bb); !bsi_end_p (bsi); bsi_next (&bsi))
        update_stmt_if_modified (bsi_stmt (bsi));
@end smallexample

@node SSA
@section Static Single Assignment
@cindex SSA
@cindex static single assignment

Most of the tree optimizers rely on the data flow information provided
by the Static Single Assignment (SSA) form.  We implement the SSA form
as described in @cite{R. Cytron, J. Ferrante, B. Rosen, M. Wegman, and
K. Zadeck.  Efficiently Computing Static Single Assignment Form and the
Control Dependence Graph.  ACM Transactions on Programming Languages
and Systems, 13(4):451-490, October 1991}.

The SSA form is based on the premise that program variables are
assigned in exactly one location in the program.  Multiple assignments
to the same variable create new versions of that variable.  Naturally,
actual programs are seldom in SSA form initially because variables
tend to be assigned multiple times.  The compiler modifies the program
representation so that every time a variable is assigned in the code,
a new version of the variable is created.  Different versions of the
same variable are distinguished by subscripting the variable name with
its version number.  Variables used in the right-hand side of
expressions are renamed so that their version number matches that of
the most recent assignment.

We represent variable versions using @code{SSA_NAME} nodes.  The
renaming process in @file{tree-ssa.c} wraps every real and
virtual operand with an @code{SSA_NAME} node which contains
the version number and the statement that created the
@code{SSA_NAME}.  Only definitions and virtual definitions may
create new @code{SSA_NAME} nodes.

Sometimes, flow of control makes it impossible to determine what is the
most recent version of a variable.  In these cases, the compiler
inserts an artificial definition for that variable called
@dfn{PHI function} or @dfn{PHI node}.  This new definition merges
all the incoming versions of the variable to create a new name
for it.  For instance,

if (...)
  a_1 = 5;
else if (...)
  a_2 = 2;
  a_3 = 13;

# a_4 = PHI <a_1, a_2, a_3>
return a_4;
@end smallexample

Since it is not possible to determine which of the three branches
will be taken at runtime, we don't know which of @code{a_1},
@code{a_2} or @code{a_3} to use at the return statement.  So, the
SSA renamer creates a new version @code{a_4} which is assigned
the result of ``merging'' @code{a_1}, @code{a_2} and @code{a_3}.
Hence, PHI nodes mean ``one of these operands.  I don't know

The following macros can be used to examine PHI nodes

@defmac	PHI_RESULT (@var{phi})
Returns the @code{SSA_NAME} created by PHI node @var{phi} (i.e.,
@var{phi}'s LHS)@.
@end defmac

@defmac	PHI_NUM_ARGS (@var{phi})
Returns the number of arguments in @var{phi}.  This number is exactly
the number of incoming edges to the basic block holding @var{phi}@.
@end defmac

@defmac	PHI_ARG_ELT (@var{phi}, @var{i})
Returns a tuple representing the @var{i}th argument of @var{phi}@.
Each element of this tuple contains an @code{SSA_NAME} @var{var} and
the incoming edge through which @var{var} flows.
@end defmac

@defmac	PHI_ARG_EDGE (@var{phi}, @var{i})
Returns the incoming edge for the @var{i}th argument of @var{phi}.
@end defmac

@defmac	PHI_ARG_DEF (@var{phi}, @var{i})
Returns the @code{SSA_NAME} for the @var{i}th argument of @var{phi}.
@end defmac

@subsection Preserving the SSA form
@findex update_ssa
@cindex preserving SSA form
Some optimization passes make changes to the function that
invalidate the SSA property.  This can happen when a pass has
added new symbols or changed the program so that variables that
were previously aliased aren't anymore.  Whenever something like this
happens, the affected symbols must be renamed into SSA form again.  
Transformations that emit new code or replicate existing statements
will also need to update the SSA form@.

Since GCC implements two different SSA forms for register and virtual
variables, keeping the SSA form up to date depends on whether you are
updating register or virtual names.  In both cases, the general idea
behind incremental SSA updates is similar: when new SSA names are
created, they typically are meant to replace other existing names in
the program@.

For instance, given the following code:

     1	L0:
     2	x_1 = PHI (0, x_5)
     3	if (x_1 < 10)
     4	  if (x_1 > 7)
     5	    y_2 = 0
     6	  else
     7	    y_3 = x_1 + x_7
     8	  endif
     9	  x_5 = x_1 + 1
     10   goto L0;
     11	endif
@end smallexample

Suppose that we insert new names @code{x_10} and @code{x_11} (lines
@code{4} and @code{8})@.

     1	L0:
     2	x_1 = PHI (0, x_5)
     3	if (x_1 < 10)
     4	  x_10 = ...
     5	  if (x_1 > 7)
     6	    y_2 = 0
     7	  else
     8	    x_11 = ...
     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

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

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

@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, any time you
have modified a statement that has virtual operands, you should verify
whether the list of virtual operands has changed, and if so, mark the
newly exposed symbols by calling @code{mark_new_vars_to_rename}.

There is one additional caveat to preserving virtual SSA form.  When the
entire set of virtual operands may be eliminated due to better
disambiguation, a bare SMT will be added to the list of virtual
operands, to signify the non-visible aliases that the are still being
referenced.  If the set of bare SMT's may change,
@code{TODO_update_smt_usage} should be added to the todo flags.

With the current pruning code, this can only occur when constants are
propagated into array references that were previously non-constant, or
address expressions are propagated into their uses.

@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:

@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:

@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:

@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.

struct foo
  int a;
  int b;
struct foo temp;
int bar (void)
  int tmp1, tmp2, tmp3;
  SFT.0_2 = V_MUST_DEF <SFT.0_1>
  temp.a = 5;
  SFT.1_4 = V_MUST_DEF <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@.
@end enumerate

For instance, consider the following function:

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

  if (i > 10)
    p = &a;
    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
Every time pointer @code{p} is dereferenced, we want to mark the
operation as a potential reference to @code{a} and @code{b}.

foo (int i)
  int *p, a, b;

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

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

  # a_9 = V_MAY_DEF <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:

@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:

may-aliases(T) = @{ V1, V2, V3, ..., 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:

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

@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:

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:

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:

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