\chapter[chap:normalization]{Normalization}
% A helper to print a single example in the half the page width. The example
% text should be in a buffer whose name is given in an argument.
%
% The align=right option really does left-alignment, but without the program
% will end up on a single line. The strut=no option prevents a bunch of empty
% space at the start of the frame.
\define[1]\example{
\framed[offset=1mm,align=right,strut=no,background=box,frame=off]{
\setuptyping[option=LAM,style=sans,before=,after=]
\typebuffer[#1]
\setuptyping[option=none,style=\tttf]
}
}
% A transformation example
\definefloat[example][examples]
\setupcaption[example][location=top] % Put captions on top
\define[3]\transexample{
\placeexample[here]{#1}
\startcombination[2*1]
{\example{#2}}{Original program}
{\example{#3}}{Transformed program}
\stopcombination
}
%
%\define[3]\transexampleh{
%% \placeexample[here]{#1}
%% \startcombination[1*2]
%% {\example{#2}}{Original program}
%% {\example{#3}}{Transformed program}
%% \stopcombination
%}
The first step in the core to \small{VHDL} translation process, is normalization. We
aim to bring the core description into a simpler form, which we can
subsequently translate into \small{VHDL} easily. This normal form is needed because
the full core language is more expressive than \small{VHDL} in some areas and because
core can describe expressions that do not have a direct hardware
interpretation.
TODO: Describe core properties not supported in \small{VHDL}, and describe how the
\small{VHDL} we want to generate should look like.
\section{Normal form}
The transformations described here have a well-defined goal: To bring the
program in a well-defined form that is directly translatable to hardware,
while fully preserving the semantics of the program. We refer to this form as
the \emph{normal form} of the program. The formal definition of this normal
form is quite simple:
\placedefinition{}{A program is in \emph{normal form} if none of the
transformations from this chapter apply.}
Of course, this is an \quote{easy} definition of the normal form, since our
program will end up in normal form automatically. The more interesting part is
to see if this normal form actually has the properties we would like it to
have.
But, before getting into more definitions and details about this normal form,
let's try to get a feeling for it first. The easiest way to do this is by
describing the things we want to not have in a normal form.
\startitemize
\item Any \emph{polymorphism} must be removed. When laying down hardware, we
can't generate any signals that can have multiple types. All types must be
completely known to generate hardware.
\item Any \emph{higher order} constructions must be removed. We can't
generate a hardware signal that contains a function, so all values,
arguments and returns values used must be first order.
\item Any complex \emph{nested scopes} must be removed. In the \small{VHDL}
description, every signal is in a single scope. Also, full expressions are
not supported everywhere (in particular port maps can only map signal names,
not expressions). To make the \small{VHDL} generation easy, all values must be bound
on the \quote{top level}.
\stopitemize
TODO: Intermezzo: functions vs plain values
A very simple example of a program in normal form is given in
\in{example}[ex:MulSum]. As you can see, all arguments to the function (which
will become input ports in the final hardware) are at the top. This means that
the body of the final lambda abstraction is never a function, but always a
plain value.
After the lambda abstractions, we see a single let expression, that binds two
variables (\lam{mul} and \lam{sum}). These variables will be signals in the
final hardware, bound to the output port of the \lam{*} and \lam{+}
components.
The final line (the \quote{return value} of the function) selects the
\lam{sum} signal to be the output port of the function. This \quote{return
value} can always only be a variable reference, never a more complex
expression.
\startbuffer[MulSum]
alu :: Bit -> Word -> Word -> Word
alu = λa.λb.λc.
let
mul = (*) a b
sum = (+) mul c
in
sum
\stopbuffer
\startuseMPgraphic{MulSum}
save a, b, c, mul, add, sum;
% I/O ports
newCircle.a(btex $a$ etex) "framed(false)";
newCircle.b(btex $b$ etex) "framed(false)";
newCircle.c(btex $c$ etex) "framed(false)";
newCircle.sum(btex $res$ etex) "framed(false)";
% Components
newCircle.mul(btex - etex);
newCircle.add(btex + etex);
a.c - b.c = (0cm, 2cm);
b.c - c.c = (0cm, 2cm);
add.c = c.c + (2cm, 0cm);
mul.c = midpoint(a.c, b.c) + (2cm, 0cm);
sum.c = add.c + (2cm, 0cm);
c.c = origin;
% Draw objects and lines
drawObj(a, b, c, mul, add, sum);
ncarc(a)(mul) "arcangle(15)";
ncarc(b)(mul) "arcangle(-15)";
ncline(c)(add);
ncline(mul)(add);
ncline(add)(sum);
\stopuseMPgraphic
\placeexample[here][ex:MulSum]{Simple architecture consisting of an adder and a
subtractor.}
\startcombination[2*1]
{\typebufferlam{MulSum}}{Core description in normal form.}
{\boxedgraphic{MulSum}}{The architecture described by the normal form.}
\stopcombination
The previous example described composing an architecture by calling other
functions (operators), resulting in a simple architecture with component and
connection. There is of course also some mechanism for choice in the normal
form. In a normal Core program, the \emph{case} expression can be used in a
few different ways to describe choice. In normal form, this is limited to a
very specific form.
\in{Example}[ex:AddSubAlu] shows an example describing a
simple \small{ALU}, which chooses between two operations based on an opcode
bit. The main structure is the same as in \in{example}[ex:MulSum], but this
time the \lam{res} variable is bound to a case expression. This case
expression scrutinizes the variable \lam{opcode} (and scrutinizing more
complex expressions is not supported). The case expression can select a
different variable based on the constructor of \lam{opcode}.
\startbuffer[AddSubAlu]
alu :: Bit -> Word -> Word -> Word
alu = λopcode.λa.λb.
let
res1 = (+) a b
res2 = (-) a b
res = case opcode of
Low -> res1
High -> res2
in
res
\stopbuffer
\startuseMPgraphic{AddSubAlu}
save opcode, a, b, add, sub, mux, res;
% I/O ports
newCircle.opcode(btex $opcode$ etex) "framed(false)";
newCircle.a(btex $a$ etex) "framed(false)";
newCircle.b(btex $b$ etex) "framed(false)";
newCircle.res(btex $res$ etex) "framed(false)";
% Components
newCircle.add(btex + etex);
newCircle.sub(btex - etex);
newMux.mux;
opcode.c - a.c = (0cm, 2cm);
add.c - a.c = (4cm, 0cm);
sub.c - b.c = (4cm, 0cm);
a.c - b.c = (0cm, 3cm);
mux.c = midpoint(add.c, sub.c) + (1.5cm, 0cm);
res.c - mux.c = (1.5cm, 0cm);
b.c = origin;
% Draw objects and lines
drawObj(opcode, a, b, res, add, sub, mux);
ncline(a)(add) "posA(e)";
ncline(b)(sub) "posA(e)";
nccurve(a)(sub) "posA(e)", "angleA(0)";
nccurve(b)(add) "posA(e)", "angleA(0)";
nccurve(add)(mux) "posB(inpa)", "angleB(0)";
nccurve(sub)(mux) "posB(inpb)", "angleB(0)";
nccurve(opcode)(mux) "posB(n)", "angleA(0)", "angleB(-90)";
ncline(mux)(res) "posA(out)";
\stopuseMPgraphic
\placeexample[here][ex:AddSubAlu]{Simple \small{ALU} supporting two operations.}
\startcombination[2*1]
{\typebufferlam{AddSubAlu}}{Core description in normal form.}
{\boxedgraphic{AddSubAlu}}{The architecture described by the normal form.}
\stopcombination
As a more complete example, consider \in{example}[ex:NormalComplete]. This
example contains everything that is supported in normal form, with the
exception of builtin higher order functions. The graphical version of the
architecture contains a slightly simplified version, since the state tuple
packing and unpacking have been left out. Instead, two seperate registers are
drawn. Also note that most synthesis tools will further optimize this
architecture by removing the multiplexers at the register input and replace
them with some logic in the clock inputs, but we want to show the architecture
as close to the description as possible.
\startbuffer[NormalComplete]
regbank :: Bit
-> Word
-> State (Word, Word)
-> (State (Word, Word), Word)
-- All arguments are an inital lambda
regbank = λa.λd.λsp.
-- There are nested let expressions at top level
let
-- Unpack the state by coercion (\eg, cast from
-- State (Word, Word) to (Word, Word))
s = sp :: (Word, Word)
-- Extract both registers from the state
r1 = case s of (fst, snd) -> fst
r2 = case s of (fst, snd) -> snd
-- Calling some other user-defined function.
d' = foo d
-- Conditional connections
out = case a of
High -> r1
Low -> r2
r1' = case a of
High -> d'
Low -> r1
r2' = case a of
High -> r2
Low -> d'
-- Packing a tuple
s' = (,) r1' r2'
-- pack the state by coercion (\eg, cast from
-- (Word, Word) to State (Word, Word))
sp' = s' :: State (Word, Word)
-- Pack our return value
res = (,) sp' out
in
-- The actual result
res
\stopbuffer
\startuseMPgraphic{NormalComplete}
save a, d, r, foo, muxr, muxout, out;
% I/O ports
newCircle.a(btex \lam{a} etex) "framed(false)";
newCircle.d(btex \lam{d} etex) "framed(false)";
newCircle.out(btex \lam{out} etex) "framed(false)";
% Components
%newCircle.add(btex + etex);
newBox.foo(btex \lam{foo} etex);
newReg.r1(btex $\lam{r1}$ etex) "dx(4mm)", "dy(6mm)";
newReg.r2(btex $\lam{r2}$ etex) "dx(4mm)", "dy(6mm)", "reflect(true)";
newMux.muxr1;
% Reflect over the vertical axis
reflectObj(muxr1)((0,0), (0,1));
newMux.muxr2;
newMux.muxout;
rotateObj(muxout)(-90);
d.c = foo.c + (0cm, 1.5cm);
a.c = (xpart r2.c + 2cm, ypart d.c - 0.5cm);
foo.c = midpoint(muxr1.c, muxr2.c) + (0cm, 2cm);
muxr1.c = r1.c + (0cm, 2cm);
muxr2.c = r2.c + (0cm, 2cm);
r2.c = r1.c + (4cm, 0cm);
r1.c = origin;
muxout.c = midpoint(r1.c, r2.c) - (0cm, 2cm);
out.c = muxout.c - (0cm, 1.5cm);
% % Draw objects and lines
drawObj(a, d, foo, r1, r2, muxr1, muxr2, muxout, out);
ncline(d)(foo);
nccurve(foo)(muxr1) "angleA(-90)", "posB(inpa)", "angleB(180)";
nccurve(foo)(muxr2) "angleA(-90)", "posB(inpb)", "angleB(0)";
nccurve(muxr1)(r1) "posA(out)", "angleA(180)", "posB(d)", "angleB(0)";
nccurve(r1)(muxr1) "posA(out)", "angleA(0)", "posB(inpb)", "angleB(180)";
nccurve(muxr2)(r2) "posA(out)", "angleA(0)", "posB(d)", "angleB(180)";
nccurve(r2)(muxr2) "posA(out)", "angleA(180)", "posB(inpa)", "angleB(0)";
nccurve(r1)(muxout) "posA(out)", "angleA(0)", "posB(inpb)", "angleB(-90)";
nccurve(r2)(muxout) "posA(out)", "angleA(180)", "posB(inpa)", "angleB(-90)";
% Connect port a
nccurve(a)(muxout) "angleA(-90)", "angleB(180)", "posB(sel)";
nccurve(a)(muxr1) "angleA(180)", "angleB(-90)", "posB(sel)";
nccurve(a)(muxr2) "angleA(180)", "angleB(-90)", "posB(sel)";
ncline(muxout)(out) "posA(out)";
\stopuseMPgraphic
\placeexample[here][ex:NormalComplete]{Simple architecture consisting of an adder and a
subtractor.}
\startcombination[2*1]
{\typebufferlam{NormalComplete}}{Core description in normal form.}
{\boxedgraphic{NormalComplete}}{The architecture described by the normal form.}
\stopcombination
\subsection{Normal form definition}
Now we have some intuition for the normal form, we can describe how we want
the normal form to look like in a slightly more formal manner. The following
EBNF-like description completely captures the intended structure (and
generates a subset of GHC's core format).
Some clauses have an expression listed in parentheses. These are conditions
that need to apply to the clause.
\startlambda
\italic{normal} = \italic{lambda}
\italic{lambda} = λvar.\italic{lambda} (representable(var))
| \italic{toplet}
\italic{toplet} = let \italic{binding} in \italic{toplet}
| letrec [\italic{binding}] in \italic{toplet}
| var (representable(varvar))
\italic{binding} = var = \italic{rhs} (representable(rhs))
-- State packing and unpacking by coercion
| var0 = var1 :: State ty (lvar(var1))
| var0 = var1 :: ty (var0 :: State ty) (lvar(var1))
\italic{rhs} = userapp
| builtinapp
-- Extractor case
| case var of C a0 ... an -> ai (lvar(var))
-- Selector case
| case var of (lvar(var))
DEFAULT -> var0 (lvar(var0))
C w0 ... wn -> resvar (\forall{}i, wi \neq resvar, lvar(resvar))
\italic{userapp} = \italic{userfunc}
| \italic{userapp} {userarg}
\italic{userfunc} = var (gvar(var))
\italic{userarg} = var (lvar(var))
\italic{builtinapp} = \italic{builtinfunc}
| \italic{builtinapp} \italic{builtinarg}
\italic{builtinfunc} = var (bvar(var))
\italic{builtinarg} = \italic{coreexpr}
\stoplambda
-- TODO: Limit builtinarg further
-- TODO: There can still be other casts around (which the code can handle,
e.g., ignore), which still need to be documented here.
-- TODO: Note about the selector case. It just supports Bit and Bool
currently, perhaps it should be generalized in the normal form?
When looking at such a program from a hardware perspective, the top level
lambda's define the input ports. The value produced by the let expression is
the output port. Most function applications bound by the let expression
define a component instantiation, where the input and output ports are mapped
to local signals or arguments. Some of the others use a builtin
construction (\eg the \lam{case} statement) or call a builtin function
(\eg \lam{add} or \lam{sub}). For these, a hardcoded \small{VHDL} translation is
available.
\subsection{Definitions}
In the following sections, we will be using a number of functions and
notations, which we will define here.
\subsubsection{Transformations}
The most important notation is the one for transformation, which looks like
the following:
\starttrans
context conditions
~
from
------------------------ expression conditions
to
~
context additions
\stoptrans
Here, we describe a transformation. The most import parts are \lam{from} and
\lam{to}, which describe the Core expresssion that should be matched and the
expression that it should be replaced with. This matching can occur anywhere
in function that is being normalized, so it applies to any subexpression as
well.
The \lam{expression conditions} list a number of conditions on the \lam{from}
expression that must hold for the transformation to apply.
Furthermore, there is some way to look into the environment (\eg, other top
level bindings). The \lam{context conditions} part specifies any number of
top level bindings that must be present for the transformation to apply.
Usually, this lists a top level binding that binds an identfier that is also
used in the \lam{from} expression, allowing us to "access" the value of a top
level binding in the \lam{to} expression (\eg, for inlining).
Finally, there is a way to influence the environment. The \lam{context
additions} part lists any number of new top level bindings that should be
added.
If there are no \lam{context conditions} or \lam{context additions}, they can
be left out alltogether, along with the separator \lam{~}.
TODO: Example
\subsubsection{Other concepts}
A \emph{global variable} is any variable that is bound at the
top level of a program, or an external module. A local variable is any other
variable (\eg, variables local to a function, which can be bound by lambda
abstractions, let expressions and case expressions).
A \emph{hardware representable} type is a type that we can generate
a signal for in hardware. For example, a bit, a vector of bits, a 32 bit
unsigned word, etc. Types that are not runtime representable notably
include (but are not limited to): Types, dictionaries, functions.
A \emph{builtin function} is a function for which a builtin
hardware translation is available, because its actual definition is not
translatable. A user-defined function is any other function.
\subsubsection{Functions}
Here, we define a number of functions that can be used below to concisely
specify conditions.
\emph{gvar(expr)} is true when \emph{expr} is a variable that references a
global variable. It is false when it references a local variable.
\emph{lvar(expr)} is the inverse of \emph{gvar}; it is true when \emph{expr}
references a local variable, false when it references a global variable.
\emph{representable(expr)} or \emph{representable(var)} is true when
\emph{expr} or \emph{var} has a type that is representable at runtime.
\section{Transform passes}
In this section we describe the actual transforms. Here we're using
the core language in a notation that resembles lambda calculus.
Each of these transforms is meant to be applied to every (sub)expression
in a program, for as long as it applies. Only when none of the
transformations can be applied anymore, the program is in normal form (by
definition). We hope to be able to prove that this form will obey all of the
constraints defined above, but this has yet to happen (though it seems likely
that it will).
Each of the transforms will be described informally first, explaining
the need for and goal of the transform. Then, a formal definition is
given, using a familiar syntax from the world of logic. Each transform
is specified as a number of conditions (above the horizontal line) and a
number of conclusions (below the horizontal line). The details of using
this notation are still a bit fuzzy, so comments are welcom.
TODO: Formally describe the "apply to every (sub)expression" in terms of
rules with full transformations in the conditions.
\subsection{Binder uniqueness}
A common problem in transformation systems, is binder uniqueness. When not
considering this problem, it is easy to create transformations that mix up
bindings and cause name collisions. Take for example, the following core
expression:
\startlambda
(λa.λb.λc. a * b * c) x c
\stoplambda
By applying β-reduction (see below) once, we can simplify this expression to:
\startlambda
(λb.λc. x * b * c) c
\stoplambda
Now, we have replaced the \lam{a} binder with a reference to the \lam{x}
binder. No harm done here. But note that we see multiple occurences of the
\lam{c} binder. The first is a binding occurence, to which the second refers.
The last, however refers to \emph{another} instance of \lam{c}, which is
bound somewhere outside of this expression. Now, if we would apply beta
reduction without taking heed of binder uniqueness, we would get:
\startlambda
λc. x * c * c
\stoplambda
This is obviously not what was supposed to happen! The root of this problem is
the reuse of binders: Identical binders can be bound in different scopes, such
that only the inner one is \quote{visible} in the inner expression. In the example
above, the \lam{c} binder was bound outside of the expression and in the inner
lambda expression. Inside that lambda expression, only the inner \lam{c} is
visible.
There are a number of ways to solve this. \small{GHC} has isolated this
problem to their binder substitution code, which performs \emph{deshadowing}
during its expression traversal. This means that any binding that shadows
another binding on a higher level is replaced by a new binder that does not
shadow any other binding. This non-shadowing invariant is enough to prevent
binder uniqueness problems in \small{GHC}.
In our transformation system, maintaining this non-shadowing invariant is
a bit harder to do (mostly due to implementation issues, the prototype doesn't
use \small{GHC}'s subsitution code). Also, we can observe the following
points.
\startitemize
\item Deshadowing does not guarantee overall uniqueness. For example, the
following (slightly contrived) expression shows the identifier \lam{x} bound in
two seperate places (and to different values), even though no shadowing
occurs.
\startlambda
(let x = 1 in x) + (let x = 2 in x)
\stoplambda
\item In our normal form (and the resulting \small{VHDL}), all binders
(signals) will end up in the same scope. To allow this, all binders within the
same function should be unique.
\item When we know that all binders in an expression are unique, moving around
or removing a subexpression will never cause any binder conflicts. If we have
some way to generate fresh binders, introducing new subexpressions will not
cause any problems either. The only way to cause conflicts is thus to
duplicate an existing subexpression.
\stopitemize
Given the above, our prototype maintains a unique binder invariant. This
meanst that in any given moment during normalization, all binders \emph{within
a single function} must be unique. To achieve this, we apply the following
technique.
TODO: Define fresh binders and unique supplies
\startitemize
\item Before starting normalization, all binders in the function are made
unique. This is done by generating a fresh binder for every binder used. This
also replaces binders that did not pose any conflict, but it does ensure that
all binders within the function are generated by the same unique supply. See
(TODO: ref fresh binder).
\item Whenever a new binder must be generated, we generate a fresh binder that
is guaranteed to be different from \emph{all binders generated so far}. This
can thus never introduce duplication and will maintain the invariant.
\item Whenever (part of) an expression is duplicated (for example when
inlining), all binders in the expression are replaced with fresh binders
(using the same method as at the start of normalization). These fresh binders
can never introduce duplication, so this will maintain the invariant.
\item Whenever we move part of an expression around within the function, there
is no need to do anything special. There is obviously no way to introduce
duplication by moving expressions around. Since we know that each of the
binders is already unique, there is no way to introduce (incorrect) shadowing
either.
\stopitemize
\subsection{η-abstraction}
This transformation makes sure that all arguments of a function-typed
expression are named, by introducing lambda expressions. When combined with
β-reduction and function inlining below, all function-typed expressions should
be lambda abstractions or global identifiers.
\starttrans
E \lam{E :: * -> *}
-------------- \lam{E} is not the first argument of an application.
λx.E x \lam{E} is not a lambda abstraction.
\lam{x} is a variable that does not occur free in \lam{E}.
\stoptrans
\startbuffer[from]
foo = λa.case a of
True -> λb.mul b b
False -> id
\stopbuffer
\startbuffer[to]
foo = λa.λx.(case a of
True -> λb.mul b b
False -> λy.id y) x
\stopbuffer
\transexample{η-abstraction}{from}{to}
\subsection{Extended β-reduction}
This transformation is meant to propagate application expressions downwards
into expressions as far as possible. In lambda calculus, this reduction
is known as β-reduction, but it is of course only defined for
applications of lambda abstractions. We extend this reduction to also
work for the rest of core (case and let expressions).
For let expressions:
\starttrans
let binds in E) M
-----------------
let binds in E M
\stoptrans
For case statements:
\starttrans
(case x of
p1 -> E1
\vdots
pn -> En) M
-----------------
case x of
p1 -> E1 M
\vdots
pn -> En M
\stoptrans
For lambda expressions:
\starttrans
(λx.E) M
-----------------
E[M/x]
\stoptrans
% And an example
\startbuffer[from]
( let a = (case x of
True -> id
False -> neg
) 1
b = (let y = 3 in add y) 2
in
(λz.add 1 z)
) 3
\stopbuffer
\startbuffer[to]
let a = case x of
True -> id 1
False -> neg 1
b = let y = 3 in add y 2
in
add 1 3
\stopbuffer
\transexample{Extended β-reduction}{from}{to}
\subsection{Let derecursification}
This transformation is meant to make lets non-recursive whenever possible.
This might allow other optimizations to do their work better. TODO: Why is
this needed exactly?
\subsection{Let flattening}
This transformation puts nested lets in the same scope, by lifting the
binding(s) of the inner let into a new let around the outer let. Eventually,
this will cause all let bindings to appear in the same scope (they will all be
in scope for the function return value).
Note that this transformation does not try to be smart when faced with
recursive lets, it will just leave the lets recursive (possibly joining a
recursive and non-recursive let into a single recursive let). The let
rederursification transformation will do this instead.
\starttrans
letnonrec x = (let bindings in M) in N
------------------------------------------
let bindings in (letnonrec x = M) in N
\stoptrans
\starttrans
letrec
\vdots
x = (let bindings in M)
\vdots
in
N
------------------------------------------
letrec
\vdots
bindings
x = M
\vdots
in
N
\stoptrans
\startbuffer[from]
let
a = letrec
x = 1
y = 2
in
x + y
in
letrec
b = let c = 3 in a + c
d = 4
in
d + b
\stopbuffer
\startbuffer[to]
letrec
x = 1
y = 2
in
let
a = x + y
in
letrec
c = 3
b = a + c
d = 4
in
d + b
\stopbuffer
\transexample{Let flattening}{from}{to}
\subsection{Empty let removal}
This transformation is simple: It removes recursive lets that have no bindings
(which usually occurs when let derecursification removes the last binding from
it).
\starttrans
letrec in M
--------------
M
\stoptrans
\subsection{Simple let binding removal}
This transformation inlines simple let bindings (\eg a = b).
This transformation is not needed to get into normal form, but makes the
resulting \small{VHDL} a lot shorter.
\starttrans
letnonrec
a = b
in
M
-----------------
M[b/a]
\stoptrans
\starttrans
letrec
\vdots
a = b
\vdots
in
M
-----------------
let
\vdots [b/a]
\vdots [b/a]
in
M[b/a]
\stoptrans
\subsection{Unused let binding removal}
This transformation removes let bindings that are never used. Usually,
the desugarer introduces some unused let bindings.
This normalization pass should really be unneeded to get into normal form
(since ununsed bindings are not forbidden by the normal form), but in practice
the desugarer or simplifier emits some unused bindings that cannot be
normalized (e.g., calls to a \type{PatError} (TODO: Check this name)). Also,
this transformation makes the resulting \small{VHDL} a lot shorter.
\starttrans
let a = E in M
---------------------------- \lam{a} does not occur free in \lam{M}
M
\stoptrans
\starttrans
letrec
\vdots
a = E
\vdots
in
M
---------------------------- \lam{a} does not occur free in \lam{M}
letrec
\vdots
\vdots
in
M
\stoptrans
\subsection{Non-representable binding inlining}
This transform inlines let bindings that have a non-representable type. Since
we can never generate a signal assignment for these bindings (we cannot
declare a signal assignment with a non-representable type, for obvious
reasons), we have no choice but to inline the binding to remove it.
If the binding is non-representable because it is a lambda abstraction, it is
likely that it will inlined into an application and β-reduction will remove
the lambda abstraction and turn it into a representable expression at the
inline site. The same holds for partial applications, which can be turned into
full applications by inlining.
Other cases of non-representable bindings we see in practice are primitive
Haskell types. In most cases, these will not result in a valid normalized
output, but then the input would have been invalid to start with. There is one
exception to this: When a builtin function is applied to a non-representable
expression, things might work out in some cases. For example, when you write a
literal \hs{SizedInt} in Haskell, like \hs{1 :: SizedInt D8}, this results in
the following core: \lam{fromInteger (smallInteger 10)}, where for example
\lam{10 :: GHC.Prim.Int\#} and \lam{smallInteger 10 :: Integer} have
non-representable types. TODO: This/these paragraph(s) should probably become a
separate discussion somewhere else.
\starttrans
letnonrec a = E in M
-------------------------- \lam{E} has a non-representable type.
M[E/a]
\stoptrans
\starttrans
letrec
\vdots
a = E
\vdots
in
M
-------------------------- \lam{E} has a non-representable type.
letrec
\vdots [E/a]
\vdots [E/a]
in
M[E/a]
\stoptrans
\startbuffer[from]
letrec
a = smallInteger 10
inc = λa -> add a 1
inc' = add 1
x = fromInteger a
in
inc (inc' x)
\stopbuffer
\startbuffer[to]
letrec
x = fromInteger (smallInteger 10)
in
(λa -> add a 1) (add 1 x)
\stopbuffer
\transexample{Let flattening}{from}{to}
\subsection{Compiler generated top level binding inlining}
TODO
\subsection{Scrutinee simplification}
This transform ensures that the scrutinee of a case expression is always
a simple variable reference.
\starttrans
case E of
alts
----------------- \lam{E} is not a local variable reference
let x = E in
case E of
alts
\stoptrans
\startbuffer[from]
case (foo a) of
True -> a
False -> b
\stopbuffer
\startbuffer[to]
let x = foo a in
case x of
True -> a
False -> b
\stopbuffer
\transexample{Let flattening}{from}{to}
\subsection{Case simplification}
This transformation ensures that all case expressions become normal form. This
means they will become one of:
\startitemize
\item An extractor case with a single alternative that picks a single field
from a datatype, \eg \lam{case x of (a, b) -> a}.
\item A selector case with multiple alternatives and only wild binders, that
makes a choice between expressions based on the constructor of another
expression, \eg \lam{case x of Low -> a; High -> b}.
\stopitemize
\starttrans
case E of
C0 v0,0 ... v0,m -> E0
\vdots
Cn vn,0 ... vn,m -> En
--------------------------------------------------- \forall i \forall j, 0 <= i <= n, 0 <= i < m (\lam{wi,j} is a wild (unused) binder)
letnonrec
v0,0 = case x of C0 v0,0 .. v0,m -> v0,0
\vdots
v0,m = case x of C0 v0,0 .. v0,m -> v0,m
x0 = E0
\dots
vn,m = case x of Cn vn,0 .. vn,m -> vn,m
xn = En
in
case E of
C0 w0,0 ... w0,m -> x0
\vdots
Cn wn,0 ... wn,m -> xn
\stoptrans
TODO: This transformation specified like this is complicated and misses
conditions to prevent looping with itself. Perhaps we should split it here for
discussion?
\startbuffer[from]
case a of
True -> add b 1
False -> add b 2
\stopbuffer
\startbuffer[to]
letnonrec
x0 = add b 1
x1 = add b 2
in
case a of
True -> x0
False -> x1
\stopbuffer
\transexample{Selector case simplification}{from}{to}
\startbuffer[from]
case a of
(,) b c -> add b c
\stopbuffer
\startbuffer[to]
letnonrec
b = case a of (,) b c -> b
c = case a of (,) b c -> c
x0 = add b c
in
case a of
(,) w0 w1 -> x0
\stopbuffer
\transexample{Extractor case simplification}{from}{to}
\subsection{Case removal}
This transform removes any case statements with a single alternative and
only wild binders.
These "useless" case statements are usually leftovers from case simplification
on extractor case (see the previous example).
\starttrans
case x of
C v0 ... vm -> E
---------------------- \lam{\forall i, 0 <= i <= m} (\lam{vi} does not occur free in E)
E
\stoptrans
\startbuffer[from]
case a of
(,) w0 w1 -> x0
\stopbuffer
\startbuffer[to]
x0
\stopbuffer
\transexample{Case removal}{from}{to}
\subsection{Argument simplification}
The transforms in this section deal with simplifying application
arguments into normal form. The goal here is to:
\startitemize
\item Make all arguments of user-defined functions (\eg, of which
we have a function body) simple variable references of a runtime
representable type. This is needed, since these applications will be turned
into component instantiations.
\item Make all arguments of builtin functions one of:
\startitemize
\item A type argument.
\item A dictionary argument.
\item A type level expression.
\item A variable reference of a runtime representable type.
\item A variable reference or partial application of a function type.
\stopitemize
\stopitemize
When looking at the arguments of a user-defined function, we can
divide them into two categories:
\startitemize
\item Arguments of a runtime representable type (\eg bits or vectors).
These arguments can be preserved in the program, since they can
be translated to input ports later on. However, since we can
only connect signals to input ports, these arguments must be
reduced to simple variables (for which signals will be
produced). This is taken care of by the argument extraction
transform.
\item Non-runtime representable typed arguments.
These arguments cannot be preserved in the program, since we
cannot represent them as input or output ports in the resulting
\small{VHDL}. To remove them, we create a specialized version of the
called function with these arguments filled in. This is done by
the argument propagation transform.
Typically, these arguments are type and dictionary arguments that are
used to make functions polymorphic. By propagating these arguments, we
are essentially doing the same which GHC does when it specializes
functions: Creating multiple variants of the same function, one for
each type for which it is used. Other common non-representable
arguments are functions, e.g. when calling a higher order function
with another function or a lambda abstraction as an argument.
The reason for doing this is similar to the reasoning provided for
the inlining of non-representable let bindings above. In fact, this
argument propagation could be viewed as a form of cross-function
inlining.
\stopitemize
TODO: Check the following itemization.
When looking at the arguments of a builtin function, we can divide them
into categories:
\startitemize
\item Arguments of a runtime representable type.
As we have seen with user-defined functions, these arguments can
always be reduced to a simple variable reference, by the
argument extraction transform. Performing this transform for
builtin functions as well, means that the translation of builtin
functions can be limited to signal references, instead of
needing to support all possible expressions.
\item Arguments of a function type.
These arguments are functions passed to higher order builtins,
like \lam{map} and \lam{foldl}. Since implementing these
functions for arbitrary function-typed expressions (\eg, lambda
expressions) is rather comlex, we reduce these arguments to
(partial applications of) global functions.
We can still support arbitrary expressions from the user code,
by creating a new global function containing that expression.
This way, we can simply replace the argument with a reference to
that new function. However, since the expression can contain any
number of free variables we also have to include partial
applications in our normal form.
This category of arguments is handled by the function extraction
transform.
\item Other unrepresentable arguments.
These arguments can take a few different forms:
\startdesc{Type arguments}
In the core language, type arguments can only take a single
form: A type wrapped in the Type constructor. Also, there is
nothing that can be done with type expressions, except for
applying functions to them, so we can simply leave type
arguments as they are.
\stopdesc
\startdesc{Dictionary arguments}
In the core language, dictionary arguments are used to find
operations operating on one of the type arguments (mostly for
finding class methods). Since we will not actually evaluatie
the function body for builtin functions and can generate
code for builtin functions by just looking at the type
arguments, these arguments can be ignored and left as they
are.
\stopdesc
\startdesc{Type level arguments}
Sometimes, we want to pass a value to a builtin function, but
we need to know the value at compile time. Additionally, the
value has an impact on the type of the function. This is
encoded using type-level values, where the actual value of the
argument is not important, but the type encodes some integer,
for example. Since the value is not important, the actual form
of the expression does not matter either and we can leave
these arguments as they are.
\stopdesc
\startdesc{Other arguments}
Technically, there is still a wide array of arguments that can
be passed, but does not fall into any of the above categories.
However, none of the supported builtin functions requires such
an argument. This leaves use with passing unsupported types to
a function, such as calling \lam{head} on a list of functions.
In these cases, it would be impossible to generate hardware
for such a function call anyway, so we can ignore these
arguments.
The only way to generate hardware for builtin functions with
arguments like these, is to expand the function call into an
equivalent core expression (\eg, expand map into a series of
function applications). But for now, we choose to simply not
support expressions like these.
\stopdesc
From the above, we can conclude that we can simply ignore these
other unrepresentable arguments and focus on the first two
categories instead.
\stopitemize
\subsubsection{Argument simplification}
This transform deals with arguments to functions that
are of a runtime representable type. It ensures that they will all become
references to global variables, or local signals in the resulting \small{VHDL}.
TODO: It seems we can map an expression to a port, not only a signal.
Perhaps this makes this transformation not needed?
TODO: Say something about dataconstructors (without arguments, like True
or False), which are variable references of a runtime representable
type, but do not result in a signal.
To reduce a complex expression to a simple variable reference, we create
a new let expression around the application, which binds the complex
expression to a new variable. The original function is then applied to
this variable.
\starttrans
M N
-------------------- \lam{N} is of a representable type
let x = N in M x \lam{N} is not a local variable reference
\stoptrans
\startbuffer[from]
add (add a 1) 1
\stopbuffer
\startbuffer[to]
let x = add a 1 in add x 1
\stopbuffer
\transexample{Argument extraction}{from}{to}
\subsubsection{Function extraction}
This transform deals with function-typed arguments to builtin functions.
Since these arguments cannot be propagated, we choose to extract them
into a new global function instead.
Any free variables occuring in the extracted arguments will become
parameters to the new global function. The original argument is replaced
with a reference to the new function, applied to any free variables from
the original argument.
This transformation is useful when applying higher order builtin functions
like \hs{map} to a lambda abstraction, for example. In this case, the code
that generates \small{VHDL} for \hs{map} only needs to handle top level functions and
partial applications, not any other expression (such as lambda abstractions or
even more complicated expressions).
\starttrans
M N \lam{M} is a (partial aplication of a) builtin function.
--------------------- \lam{f0 ... fn} = free local variables of \lam{N}
M x f0 ... fn \lam{N :: a -> b}
~ \lam{N} is not a (partial application of) a top level function
x = λf0 ... λfn.N
\stoptrans
\startbuffer[from]
map (λa . add a b) xs
map (add b) ys
\stopbuffer
\startbuffer[to]
x0 = λb.λa.add a b
~
map x0 xs
x1 = λb.add b
map x1 ys
\stopbuffer
\transexample{Function extraction}{from}{to}
\subsubsection{Argument propagation}
This transform deals with arguments to user-defined functions that are
not representable at runtime. This means these arguments cannot be
preserved in the final form and most be {\em propagated}.
Propagation means to create a specialized version of the called
function, with the propagated argument already filled in. As a simple
example, in the following program:
\startlambda
f = λa.λb.a + b
inc = λa.f a 1
\stoplambda
we could {\em propagate} the constant argument 1, with the following
result:
\startlambda
f' = λa.a + 1
inc = λa.f' a
\stoplambda
Special care must be taken when the to-be-propagated expression has any
free variables. If this is the case, the original argument should not be
removed alltogether, but replaced by all the free variables of the
expression. In this way, the original expression can still be evaluated
inside the new function. Also, this brings us closer to our goal: All
these free variables will be simple variable references.
To prevent us from propagating the same argument over and over, a simple
local variable reference is not propagated (since is has exactly one
free variable, itself, we would only replace that argument with itself).
This shows that any free local variables that are not runtime representable
cannot be brought into normal form by this transform. We rely on an
inlining transformation to replace such a variable with an expression we
can propagate again.
\starttrans
x = E
~
x Y0 ... Yi ... Yn \lam{Yi} is not of a runtime representable type
--------------------------------------------- \lam{Yi} is not a local variable reference
x' y0 ... yi-1 f0 ... fm Yi+1 ... Yn \lam{f0 ... fm} = free local vars of \lam{Yi}
~
x' = λy0 ... yi-1 f0 ... fm yi+1 ... yn .
E y0 ... yi-1 Yi yi+1 ... yn
\stoptrans
TODO: Example
\subsection{Cast propagation / simplification}
This transform pushes casts down into the expression as far as possible. Since
its exact role and need is not clear yet, this transformation is not yet
specified.
\subsection{Return value simplification}
This transformation ensures that the return value of a function is always a
simple local variable reference.
Currently implemented using lambda simplification, let simplification, and
top simplification. Should change into something like the following, which
works only on the result of a function instead of any subexpression. This is
achieved by the contexts, like \lam{x = E}, though this is strictly not
correct (you could read this as "if there is any function \lam{x} that binds
\lam{E}, any \lam{E} can be transformed, while we only mean the \lam{E} that
is bound by \lam{x}. This might need some extra notes or something).
\starttrans
x = E \lam{E} is representable
~ \lam{E} is not a lambda abstraction
E \lam{E} is not a let expression
--------------------------- \lam{E} is not a local variable reference
let x = E in x
\stoptrans
\starttrans
x = λv0 ... λvn.E
~ \lam{E} is representable
E \lam{E} is not a let expression
--------------------------- \lam{E} is not a local variable reference
let x = E in x
\stoptrans
\starttrans
x = λv0 ... λvn.let ... in E
~ \lam{E} is representable
E \lam{E} is not a local variable reference
---------------------------
let x = E in x
\stoptrans
\startbuffer[from]
x = add 1 2
\stopbuffer
\startbuffer[to]
x = let x = add 1 2 in x
\stopbuffer
\transexample{Return value simplification}{from}{to}