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Rewrite body text of higher order translation for deep handlers.

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Daniel Hillerström
2020-09-22 15:57:26 +01:00
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@@ -2453,7 +2453,7 @@ Since we now use a list representation for the stacks of
continuations, we have had to modify the translations of all the continuations, we have had to modify the translations of all the
constructs that manipulate continuations. For $\Return$ and $\Let$, we constructs that manipulate continuations. For $\Return$ and $\Let$, we
extract the top continuation $k$ and manipulate it analogously to the extract the top continuation $k$ and manipulate it analogously to the
original translation in Figure\ \ref{fig:cps-cbv}. For $\Do$, we original translation in Figure~\ref{fig:cps-cbv}. For $\Do$, we
extract the top pure continuation $k$ and effect continuation $h$ and extract the top pure continuation $k$ and effect continuation $h$ and
invoke $h$ in the same way as the curried translation, except that we invoke $h$ in the same way as the curried translation, except that we
explicitly maintain the stack $ks$ of additional continuations. The explicitly maintain the stack $ks$ of additional continuations. The
@@ -2765,57 +2765,74 @@ Figure~\ref{fig:cps-higher-order-uncurried}. In essence, it is the
same as the refined first-order uncurried CPS translation, although same as the refined first-order uncurried CPS translation, although
the notation is slightly more involved due to the separation of static the notation is slightly more involved due to the separation of static
and dynamic parts. and dynamic parts.
%
The translation on values is almost homomorphic on the syntax As before, the translation comprises three translation functions, one
constructors. The notable exceptions are the translations of for each syntactic category: values, computations, and handler
$\lambda$-abstractions and $\Lambda$-abstractions. The former gives definitions. Amongst the three functions, the translation function for
rise to dynamic computation at runtime, and thus, the translation computations stands out, because it is the only one that operates on
emits a dynamic $\dlam$-abstraction with two formal parameters: the static continuations. Its type signature,
function parameter, $x$, and the continuation parameter, $ks$. The $\cps{-} : \CompCat \to \SValCat^\ast \to \UCompCat$, signifies that
body of the $\dlam$-abstraction is somewhat anomalous as the it is a binary function, taking a $\HCalc$-computation term as its
continuation $ks$ is subject to immediate dynamic deconstruction just first argument and a static continuation (a list of static values) as
to be reconstructed again and used as a static argument to the its second argument, and ultimately produces a $\UCalc$-computation
translation of the original body computation $M$. This deconstruction term. Thus the computation translation function is able to manipulate
and reconstruction may at first glance appear to be rather pointless, the continuation. In fact, the translation is said to be higher-order
however, a closer look reveals that it performs a critical role as it because the continuation parameter is a higher-order: it is a list of
ensures that the translation on computation terms has immediate access functions.
to the next pure and effect continuation functions. In fact, the
translation is set up such that every generated dynamic
$\dlam$-abstraction deconstructs its continuation argument
appropriately to expose just enough (static) continuation structure in
order to guarantee that static pattern matching never fails.
%
(The translation is somewhat unsatisfactory regarding this
deconstruction of dynamic continuations as the various generated
dynamic lambda abstractions do not deconstruct their continuation
arguments uniformly.)
%
Since the target language is oblivious to types any
$\Lambda$-abstraction gets translated in a similar way to
$\lambda$-abstraction, where the first parameter becomes a placeholder
for a ``dummy'' unit argument.
The translation of computation terms is more interesting. Note the To ensure that static continuation manipulation is well-defined the
type of the translation function, translation maintains the invariant that the statically known
$\cps{-} : \CompCat \to \SValCat^\ast \to \UCompCat$. It is a binary continuation stack ($\sk$) always contains at least two continuation
function, taking a $\HCalc$-computation term as its first argument and functions, i.e. a complete continuation pair consisting of a pure
a static continuation as its second argument, and ultimately produces continuation function and an effect continuation function.
a $\UCalc$-computation term. The static continuation may be %
manipulated during the translation as is done in the translations of This invariant guarantees that all translations are uniform in whether
$\Return$, $\Let$, $\Do$, and $\Handle$. The translation of $\Return$ they appear statically within the scope of a handler or not, and this
consumes the next pure continuation function $\sk$ and applies it also simplifies the correctness proof
dynamically to the translation of $V$ and the reified remainder, (Theorem~\ref{thm:ho-simulation}).
$\sks$, of the static continuation. Keep in mind that the translation %
of $\Return$ only consumes one of the two guaranteed available Maintaining this invariant has a cosmetic effect on the presentation
continuation functions from the provided continuation argument; as a of the translation. This effect manifests in any place where a
consequence the next continuation function in $\sks$ is an effect dynamically known continuation stack is passed in (as a continuation
continuation. This consequence is accounted for in the translations of parameter $\dhk$), as it must be deconstructed using a dynamic
$\Let$ and $\Return$-clauses, which are precisely the two possible language $\Let$ to expose the continuation structure and subsequently
(dynamic) origins of $\sk$. For instance, the translation of $\Let$ reconstructed as a static value with reflected variable names.
The translation of $\lambda$-abstractions provides an example of this
deconstruction and reconstruction in action. The dynamic continuation
$\dhk$ is deconstructed to expose to the next pure continuation
function $\dk$ and effect continuation $h$, and the remainder of the
continuation $\dhk'$; these names are immediately reflected and put
back together to form a static continuation that is provided to the
translation of the body computation $M$.
The only translation rule that consumes a complete reflected
continuation pair is the translation of $\Do$. The effect continuation
function, $\sh$, is dynamically applied to an operation package and
the reified remainder of the continuation $\sks$. As usual, the
operation package contains the payload and the resumption, which is
represented as a reversed continuation slice.
%
The only other translation rules that manipulate the continuation are
$\Return$ and $\Let$, which only consume the pure continuation
function $\sk$. For example, the translation of $\Return$ is a dynamic
application of $\sk$ to the translation of the value $V$ and the
remainder of the continuation $\sks$.
%
The shape of $\sks$ is odd, meaning that the top element is an effect
continuation function. Thus the pure continuation $\sk$ has to account
for this odd shape. Fortunately, the possible instantiations of the
pure continuation are few. We can derive the all possible
instantiations systematically by using the operational semantics of
$\HCalc$. According to the operational semantics the continuation of a
$\Return$-computation is either the continuation of a
$\Let$-expression or a $\Return$-clause.
%
The translations of $\Let$-expressions and $\Return$-clauses each
account for odd continuations. For example, the translation of $\Let$
consumes the current pure continuation function and generates a consumes the current pure continuation function and generates a
replacement: a pure continuation function which expects an odd dynamic replacement: a pure continuation function which expects an odd dynamic
continuation $ks$, which it deconstructs to expose the effect continuation $\dhk$, which it deconstructs to expose the effect
continuation $h$ along with the current pure continuation function in continuation $h$ along with the current pure continuation function in
the translation of $N$. The modified continuation is passed to the the translation of $N$. The modified continuation is passed to the
translation of $M$. translation of $M$.
@@ -2880,32 +2897,169 @@ continuation in order to bind the current effect continuation function
to the name $h$, which would have been used during the translation of to the name $h$, which would have been used during the translation of
$N$. $N$.
The translation of $\Do$ consumes both the current pure continuation The translation of $\Handle$ applies the translation of $M$ to the
function and current effect continuation function. It applies the current continuation extended with the translation of the
effect continuation function dynamically to an operation package and $\Return$-clause, acting as a pure continuation function, and the
the reified remainder of the continuation. As usual, the operation translation of operation-clauses, acting as an effect continuation
package contains the payload and the resumption, which is a reversed function.
slice of the provided continuation. The translation of $\Handle$
applies the translation of $M$ to the current continuation extended
with the translation of the $\Return$-clause, acting as a pure
continuation function, and the translation of operation-clauses,
acting as an effect continuation function.
The translation of a return clause $\hret$ of some handler $H$
generates a dynamic lambda abstraction, which expects an odd
continuation. In its body it extracts three elements from the
continuation parameter $ks$. The first element is discarded as it is
the effect continuation corresponding to the translation of
$\hops$. The other two elements are used to statically expose the next
pure and effect continuation functions in the translation of $N$.
% %
The translation of $\hops$ also generates a dynamic lambda The translation of a $\Return$-clause discards the effect continuation
abstraction, which unpacks the operation package to perform a $h$ and in addition exposes the next pure continuation $\dk$ and
case-split on the operation label $z$. The continuation $ks$ is effect continuation $h'$ which are reflected to form a static
deconstructed in the branch for $\ell$ in order to expose the continuation for the translation of $N$.
%
The translation of operation clauses unpacks the provided operation
package to perform a case-split on the operation label $z$. The branch
for $\ell$ deconstructs the continuation $\dhk$ in order to expose the
continuation structure. The forwarding branch also deconstructs the continuation structure. The forwarding branch also deconstructs the
continuation, however, for a different purpose, namely, to augment the continuation, but for a different purpose; it augments the resumption
resumption $rs$ the next pure and effect continuation functions. $\dhkr$ with the next pure and effect continuation functions.
% The translation of
% $\Handle$ applies the translation of $M$ to the current continuation
% extended with the translation of the $\Return$-clause, acting as a
% pure continuation function, and the translation of operation-clauses,
% acting as an effect continuation function.
% The translation on values is almost homomorphic on the syntax
% constructors. The notable exceptions are the translations of
% $\lambda$-abstractions and $\Lambda$-abstractions. The former gives
% rise to dynamic computation at runtime, and thus, the translation
% emits a dynamic $\dlam$-abstraction with two formal parameters: the
% function parameter, $x$, and the continuation parameter, $ks$. The
% body of the $\dlam$-abstraction is somewhat anomalous as the
% continuation $ks$ is subject to immediate dynamic deconstruction just
% to be reconstructed again and used as a static argument to the
% translation of the original body computation $M$. This deconstruction
% and reconstruction may at first glance appear to be rather pointless,
% however, a closer look reveals that it performs a critical role as it
% ensures that the translation on computation terms has immediate access
% to the next pure and effect continuation functions. In fact, the
% translation is set up such that every generated dynamic
% $\dlam$-abstraction deconstructs its continuation argument
% appropriately to expose just enough (static) continuation structure in
% order to guarantee that static pattern matching never fails.
% %
% (The translation is somewhat unsatisfactory regarding this
% deconstruction of dynamic continuations as the various generated
% dynamic lambda abstractions do not deconstruct their continuation
% arguments uniformly.)
% %
% Since the target language is oblivious to types any
% $\Lambda$-abstraction gets translated in a similar way to
% $\lambda$-abstraction, where the first parameter becomes a placeholder
% for a ``dummy'' unit argument.
% The translation of computation terms is more interesting. Note the
% type of the translation function,
% $\cps{-} : \CompCat \to \SValCat^\ast \to \UCompCat$. It is a binary
% function, taking a $\HCalc$-computation term as its first argument and
% a static continuation as its second argument, and ultimately produces
% a $\UCalc$-computation term. The static continuation may be
% manipulated during the translation as is done in the translations of
% $\Return$, $\Let$, $\Do$, and $\Handle$. The translation of $\Return$
% consumes the next pure continuation function $\sk$ and applies it
% dynamically to the translation of $V$ and the reified remainder,
% $\sks$, of the static continuation. Keep in mind that the translation
% of $\Return$ only consumes one of the two guaranteed available
% continuation functions from the provided continuation argument; as a
% consequence the next continuation function in $\sks$ is an effect
% continuation. This consequence is accounted for in the translations of
% $\Let$ and $\Return$-clauses, which are precisely the two possible
% (dynamic) origins of $\sk$. For instance, the translation of $\Let$
% consumes the current pure continuation function and generates a
% replacement: a pure continuation function which expects an odd dynamic
% continuation $ks$, which it deconstructs to expose the effect
% continuation $h$ along with the current pure continuation function in
% the translation of $N$. The modified continuation is passed to the
% translation of $M$.
% %
% To provide a flavour of how this continuation manipulation functions
% in practice, consider the following example term.
% %
% \begin{derivation}
% &\pcps{\Let\;x \revto \Return\;V\;\In\;N}\\
% =& \reason{definition of $\pcps{-}$}\\
% &\ba[t]{@{}l}(\slam \sk \scons \sks.\cps{\Return\;V} \sapp
% (\reflect(\dlam x\,ks.
% \ba[t]{@{}l}
% \Let\;(h \dcons ks') = ks \;\In\\
% \cps{N} \sapp (\sk \scons \reflect h \scons \reflect ks')) \scons \sks)
% \ea\\
% \sapp (\reflect (\dlam x\,ks.x) \scons \reflect (\dlam z\,ks.\Absurd~z) \scons \snil))
% \ea\\
% =& \reason{definition of $\cps{-}$}\\
% &\ba[t]{@{}l}(\slam \sk \scons \sks.(\slam \sk \scons \sks. \reify \sk \dapp \cps{V} \dapp \reify \sks) \sapp
% (\reflect(\dlam x\,ks.
% \ba[t]{@{}l}
% \Let\;(h \dcons ks') = ks \;\In\\
% \cps{N} \sapp (\sk \scons \reflect h \scons \reflect ks')) \scons \sks)
% \ea\\
% \sapp (\reflect (\dlam x\,ks.x) \scons \reflect (\dlam z\,ks.\Absurd~z) \scons \snil))
% \ea\\
% =& \reason{static $\beta$-reduction}\\
% &(\slam \sk \scons \sks. \reify \sk \dapp \cps{V} \dapp \reify \sks)
% \sapp
% (\reflect(\dlam x\,\dhk.
% \ba[t]{@{}l}
% \Let\;(h \dcons \dhk') = \dhk \;\In\\
% \cps{N} \sapp
% \ba[t]{@{}l}
% (\reflect (\dlam x\,\dhk.x) \scons \reflect h \scons \reflect \dhk'))\\
% ~~\scons \reflect (\dlam z\,\dhk.\Absurd~z) \scons \snil))
% \ea
% \ea\\
% =& \reason{static $\beta$-reduction}\\
% &\ba[t]{@{}l@{~}l}
% &(\dlam x\,\dhk.
% \Let\;(h \dcons \dhk') = \dhk \;\In\;
% \cps{N} \sapp
% (\reflect (\dlam x\,\dhk.x) \scons \reflect h \scons \reflect \dhk'))\\
% \dapp& \cps{V} \dapp ((\dlam z\,\dhk.\Absurd~z) \dcons \dnil)\\
% \ea\\
% \reducesto& \reason{\usemlab{App_2}}\\
% &\Let\;(h \dcons \dhk') = (\dlam z\,\dhk.\Absurd~z) \dcons \dnil \;\In\;
% \cps{N[V/x]} \sapp
% (\reflect (\dlam x\,\dhk.x) \scons \reflect h \scons \reflect \dhk'))\\
% \reducesto^+& \reason{dynamic pattern matching and substitution}\\
% &\cps{N[V/x]} \sapp
% (\reflect (\dlam x\,\dhk.x) \scons \reflect (\dlam z\,\dhk.\Absurd~z) \scons \reflect \dnil)
% \end{derivation}
% %
% The translation of $\Return$ provides the generated dynamic pure
% continuation function with the odd continuation
% $((\dlam z\,ks.\Absurd~z) \dcons \dnil)$. After the \usemlab{App_2}
% reduction, the pure continuation function deconstructs the odd
% continuation in order to bind the current effect continuation function
% to the name $h$, which would have been used during the translation of
% $N$.
% The translation of $\Do$ consumes both the current pure continuation
% function and current effect continuation function. It applies the
% effect continuation function dynamically to an operation package and
% the reified remainder of the continuation. As usual, the operation
% package contains the payload and the resumption, which is a reversed
% slice of the provided continuation. The translation of $\Handle$
% applies the translation of $M$ to the current continuation extended
% with the translation of the $\Return$-clause, acting as a pure
% continuation function, and the translation of operation-clauses,
% acting as an effect continuation function.
% The translation of a return clause $\hret$ of some handler $H$
% generates a dynamic lambda abstraction, which expects an odd
% continuation. In its body it extracts three elements from the
% continuation parameter $ks$. The first element is discarded as it is
% the effect continuation corresponding to the translation of
% $\hops$. The other two elements are used to statically expose the next
% pure and effect continuation functions in the translation of $N$.
% %
% The translation of $\hops$ also generates a dynamic lambda
% abstraction, which unpacks the operation package to perform a
% case-split on the operation label $z$. The continuation $ks$ is
% deconstructed in the branch for $\ell$ in order to expose the
% continuation structure. The forwarding branch also deconstructs the
% continuation, however, for a different purpose, namely, to augment the
% resumption $rs$ the next pure and effect continuation functions.
% %
\dhil{Remark that an alternative to cluttering the translations with \dhil{Remark that an alternative to cluttering the translations with
unpacking and repacking of continuations would be to make static unpacking and repacking of continuations would be to make static
@@ -2913,75 +3067,6 @@ resumption $rs$ the next pure and effect continuation functions.
may reduce the dynamic administrative reductions, it makes the may reduce the dynamic administrative reductions, it makes the
simulation proof more complicated.} simulation proof more complicated.}
% : the $\dlam$-abstractions generated by translating $\lambda$-abstractions, $\Lambda$-abstractions, and operation clauses expose the next two continuation functions, whilst the dynamic lambda abstraction generated by translating $\Let$ only expose the next continuation, others expose only the next one, and the translation of
% $\lambda$-abstractions expose the next two continuation functions,
% The translation on values and handler definitions generates dynamic
% lambda abstractions. Whilst the translation on computations generates
% static lambda abstractions.
% Static continuation application corresponds to
% The translation function for computation terms is staged: for any
% given computation term the translation yields a static function, which
% can be applied to a static continuation to ultimately yield a
% $\UCalc$-computation term. Staging is key to eliminating static
% administrative redexes as any static function may perform static
% computation by manipulating its provided static continuation
% a binary higher-order
% function; it is higher-order because its second parameter is a list of
% static continuation functions.
% A major difference that has a large cosmetic effect on the
% presentation of the translation is that we maintain the invariant that
% the statically known stack ($\sk$) always contains at least one frame,
% consisting of a triple
% $\sRecord{\reflect V_{fs}, \sRecord{\reflect V_{ret}, \reflect
% V_{ops}}}$ of reflected dynamic pure frame stacks, return
% handlers, and operation handlers. Maintaining this invariant ensures
% that all translations are uniform in whether they appear statically
% within the scope of a handler or not, and this simplifies our
% correctness proof. To maintain the invariant, any place where a
% dynamically known stack is passed in (as a continuation parameter
% $k$), it is immediately decomposed using a dynamic language $\Let$ and
% repackaged as a static value with reflected variable
% names. Unfortunately, this does add some clutter to the translation
% definition, as compared to the translations above. However, there is a
% payoff in the removal of administrative reductions at run time. The
% translations presented by \citet{HillerstromLAS17} and
% \citet{HillerstromL18} did not do this decomposition and repackaging
% step, which resulted in additional administrative reductions in the
% translation due to the translations of $\Let$ and $\Do$ being passed
% dynamic continuations when they were expecting statically known ones.
%
% The translation on values is mostly homomorphic on the syntax
% constructors, except for $\lambda$-abstractions and
% $\Lambda$-abstractions which are translated in the exact same way,
% because the target calculus has no notion of types. As per usual
% continuation passing style practice, the translation of a lambda
% abstraction yields a dynamic binary lambda abstraction, where the
% second parameter is intended to be the continuation
% parameter.
% %
% However, the translation slightly diverge from the usual
% practice in the translation of the body as the result of $\cps{M}$ is
% applied statically, rather than dynamically, to the (reflected)
% continuation parameter.
% %
% The reason is that the translation on computations is staged: for any
% given computation term the translation yields a static function, which
% can be applied to a static continuation to ultimately produce a
% $\UCalc$-computation term. This staging is key to eliminating static
% administrative redexes as any static function may perform static
% computation by manipulating its provided static continuation, as is
% for instance done in the translation of $\Let$.
% %
\dhil{Spell out why this translation qualifies as `higher-order'.} \dhil{Spell out why this translation qualifies as `higher-order'.}
@@ -3120,7 +3205,7 @@ reductions, i.e.
% %
\[ \[
\cps{M} \sapp (\sV_1 \scons \dots \sV_n \scons \reflect \reify \sW) \cps{M} \sapp (\sV_1 \scons \dots \sV_n \scons \reflect \reify \sW)
\areducesto^* \cps{M} \sapp (\sV_1 \scons \dots \sV_n \scons \sW) \areducesto^\ast \cps{M} \sapp (\sV_1 \scons \dots \sV_n \scons \sW)
\] \]
\end{lemma} \end{lemma}
% %
@@ -3406,53 +3491,6 @@ interpretation of continuations.
Although it is seductive to program with lists, it quickly gets Although it is seductive to program with lists, it quickly gets
unwieldy. unwieldy.
% The underlying issue is that the continuation structure is still too
% extensional. The need for a more intensional structure is evident in
% the insertion of $\kid$ in the translation of $\Handle$ as a
% placeholder for the empty pure continuation. Another telltale for a
% more general notion of continuation is that continuation
% deconstructions are non-uniform due to the single flat list
% continuation representation. While stuffing everything into lists is
% seductive, it quickly gets unwieldy.
% Instead of inserting dummy handlers, one may consider composing the
% current pure continuation with the next pure continuation, meaning
% that the current pure continuation would be handled accordingly by the
% next handler. To retrieve the next pure continuation, we must reach
% under the next handler, i.e. something along the lines of the
% following.
% %
% \[
% \bl
% \Let\;(\_ \dcons \_ \dcons \dk \dcons \vhops \dcons \vhret \dcons \dk' \dcons rs') = rs\;\In\\
% \Let\;r = \Res\,(\vhops \dcons \vhret \dcons (\dk' \circ \dk) \dcons rs')\;\In\;\dots
% \el
% \]
% %
% where $\circ$ is defined to be function composition in continuation
% passing style.
% %
% \[
% g \circ f \defas \lambda x\,\dhk.
% \ba[t]{@{}l}
% \Let\;(\dk \dcons \dhk') = \dhk\; \In\\
% f \dapp x \dapp ((\lambda x\,\dhk. g \dapp x \dapp (\dk \dcons \dhk)) \dcons \dhk')
% \ea
% \]
% %
% This idea is cute, but it reintroduces a variation of the dynamic
% redexes we dealt with in
% Section~\ref{sec:first-order-explicit-resump}.
% %
% Both solutions side step the underlying issue, namely, that the
% continuation structure is still too extensional. The need for a more
% intensional structure is evident in the insertion of $\kid$ in the
% translation of $\Handle$ as a placeholder for the empty pure
% continuation. Another telltale is that the disparity of continuation
% deconstructions also gets bigger. The continuation representation
% needs more structure. While it is seductive to program with lists, it
% quickly gets unwieldy.
\subsection{Generalised continuations} \subsection{Generalised continuations}
\label{sec:generalised-continuations} \label{sec:generalised-continuations}