run
and cross-stage persistence[The Abstract of the paper]
Generating high-performance code and applying typical optimizations
within the bodies of loops and functions involves moving or storing
open code for later use, often in a different binding
environment. There are ample opportunities for variables being left
unbound or accidentally captured. It has been a tough challenge to
statically ensure that by construction the generated code is
nevertheless well-typed and well-scoped: all free variables in
manipulated and stored code fragments shall eventually be bound, by
their intended binders.
We present the calculus for code generation with mutable state that for the first time achieves type-safety and hygiene without ad hoc restrictions. The calculus strongly resembles region-based memory management, but with the orders of magnitude simpler proofs. It employs the rightly abstract representation for free variables, which, like hypothesis in natural deduction, are free from the bureaucracy of syntax imposed by the type environment or numbering conventions.
Although the calculus was designed for the sake of formalization and is deliberately bare-bone, it turns out easily implementable and not too bothersome for writing realistic program.
Joint work with Yukiyoshi Kameyama and Yuto Sudo.
APLAS-talk.pdf [294K]
Annotated slides of the talk presented at APLAS 2016 on November 22, 2016. Hanoi, Vietnam.
metaNJ.ml [14K]
Embedding of the <NJ> calculus in OCaml, with many examples
Not just any solution is acceptable: we wish to avoid tree hacking, modification or even examination of the generated code, and its post-validation. See the slide notes for the detailed explanation of these requirements. Our goal is to generate code with compositional combinators that statically assure the results (even intermediate, open results) are well-formed and well-typed.
Template Haskell is far away from the goal. MetaOCaml is closer: generators without side-effects satisfy our requirements. Alas, such generators cannot implement our benchmarks. Without effects, let-insertion or other movement of open code past generated binders is not possible.
Granted, let-insertion without crossing binders can be done effectlessly,
as well-known from partial evaluation. The cost is
writing generators in the continuation-passing, or monadic style (which
obscures the algorithm and makes the generators harder to use by domain
experts). However, even the repeated continuation-passing transformation
cannot help us insert let
beyond the closest binder.
We need a new CPS hierarchy.
Joint work with Chung-chieh Shan and Yukiyoshi Kameyama.
talk-problems.pdf [172K]
Computational Effects across Generated Binders. Part 1: Problems and solutions
Extensively annotated slides of the talk presented at the IFIP WG2.11 meeting (Bordeaux, France, September 5, 2011) and at INRIA Paris (September 9, 2011).
talk-problems.ml [11K]
MetaOCaml code for the code generation benchmarks that emphasize effects crossing future-stage binders. Although the code implements all benchmarks, the absence of scope extrusion cannot be assured. Small mistakes can indeed result in unbound variables in the generated code.
Generating optimal stencil code
A real-life example of code generation with effects crossing many binders
This paper takes a first step towards solving the problem, by translating the staging away. Our source language models MetaOCaml restricted to one future stage. It is a call-by-value language, with a sound type system and a small-step operational semantics, that supports building open code, running closed code, cross-stage persistence, and non-termination effects. We translate each typing derivation from this source language to the unstaged System F with constants. Our translation represents future-stage code using closures, yet preserves the typing, alpha-equivalence (hygiene), and (we conjecture) termination and evaluation order of the staged program.
To decouple evaluation from scope (a defining characteristic of staging), our translation weakens the typing environment of open code using a term coercion reminiscent of Goedel's translation from intuitionistic to modal logic. By converting open code to closures with typed environments, our translation establishes a framework in which to study staging with effects and to prototype staged languages. It already makes scope extrusion a type error.
Joint work with Yukiyoshi Kameyama and Chung-chieh Shan.
Chung-chieh Shan. Slides of the talk at PEPM, January 7, 2008.
< http://www.cs.rutgers.edu/~ccshan/metafx/talk.pdf >
paper-examples.ml [11K]
Examples of the staged code and its translation. The file contains
the complete code for all the examples in the paper plus a few extra.
The source language is MetaOCaml. The target language is supposed to be
System F. We try to emulate it in plain OCaml, using first-class (record)
polymorphism where needed.
power-count.ml [4K]
Computing a staged power function while tracking the number
of multiplications: The example in Sec 6 of the paper.
It is easiest to write this example with a side effect such as
mutable state in MetaOCaml, but such an extension (a piece of state of
type int
) has not been shown sound except through our translation.
Second, we can write this example in pure MetaOCaml (more awkwardly) using
our environment-passing translation.
Two-level staged calculus with environment classifiers, run
and cross-stage persistence
The source language of the translation
Two pillars of this approach are types and effects. Typed multilevel languages such as MetaOCaml ensure safety: a well-typed code generator neither goes wrong nor generates code that goes wrong. Side effects such as state and control ease correctness: an effectful generator can resemble the textbook presentation of an algorithm, as is familiar to domain experts, yet insert let
for memoization and if
for bounds-checking, as is necessary for efficiency. However, adding effects blindly renders multilevel types unsound.
We introduce the first multilevel calculus with control effects and a sound type system. We give small-step operational semantics as well as a one-pass continuation-passing-style translation. For soundness, our calculus restricts the code generator's effects to the scope of generated binders. Even with this restriction, we can finally write efficient code generators for dynamic programming and numerical methods in direct style, like in algorithm textbooks, rather than in continuation-passing or monadic style.
Joint work with Yukiyoshi Kameyama and Chung-chieh Shan.
Chung-chieh Shan. Slides of the talk at PEPM, January 20, 2009.
< http://www.cs.rutgers.edu/~ccshan/metafx/meta-shift-talk.pdf >
circle-shift.elf [38K]
lambda-circle calculus with shift/reset: two-stage calculus with delimited control effects
This Twelf code defines the calculus and its static (type checking) and dynamic semantics. The code contains many examples, including the staged Gibonacci example with memoization and let-insertion.
Mechanizing multilevel metatheory with control effects
Detailed description of the formalization of the extended calculus, with arbitrarily many levels
fib.ml [16K]
fib1.ml [11K]
The Gibonacci example -- generating efficient specialized versions of the generalized Fibonacci function in direct style
The two MetaOCaml files describe a progression of attempts, from an inefficient unstaged Gibonacci, memoized Gibonacci, naively staged and inefficient function and finally to the efficient memoization with let-insertion. The dangers of scope extrusion are well illustrated. The specialized Gibonacci generator is written in direct, rather than monadic or CPS style. The file fib.ml
uses the explicit fix-point combinator; The other file relies on recursive definitions instead.
lcs.ml [6K]
The complete MetaOCaml code for generating optimal specialized code for the longest common subsequence: another example of dynamic meta-programming in direct style
ge_unstaged.ml [8K]
ge_gen.ml [15K]
Generating a family of Gaussian Elimination codes in direct style, without either functors or monads
The file ge_unstaged.ml
is the unstaged, textbook Gaussian elimination code, in plain OCaml. The other file is the corresponding staged code, in MetaOCaml. The staged code generates a family of GE codes (with or without determinant computation, etc). We pay no performance penalty for the added flexibility.
We embedded this type system in a Haskell library. We used this tagless-staged library to implement statically safe let-insertion across an arbitrary number of binders for the first time.
Joint work with Chung-chieh Shan and Yukiyoshi Kameyama.
Haskell code generation library that statically enforces future-stage lexical scope
Unsafe.hs [19K]
Demonstrating run-time errors that our type system statically prevents. The code shows that well-scoped De Bruijn indices do not statically determine lexical scope.
We present a new translation that maps a quoted let-binding to a metalanguage let-binding, which is generalizable. It works with the current OCaml, demonstrating the benefits of the relaxed value restriction for the other ML languages. It solves the problem in practice, and almost in theory: It relies on a conjecture, whose proof is a research program, promising to re-open the old value-restriction wounds and heal them.
polylet.ml [17K]
OCaml code accompanying the paper
Some part of the code relies on the delimcc library of delimited control.
Staged let-generalization may be unsound The exception to the generalization rule proposed in the paper. We conjecture this is the only exception.
The distinguished feature of Hindley-Milner type system is generalization of let-bindings. The type inferred for a binding introduced by the let-form is generalized by quantifying generalizable free type variables. For example, the expression
let f () = [] in (1::f(), "123"::f())
is well-typed: since the type inferred for f
, unit -> 'a list
,
contains the generalizable type variable 'a
the type is
generalized to the polymorphic forall 'a. unit -> 'a list
(in
Hindley-Milner systems, quantifiers are often omitted). The
(implicitly) quantified type variable 'a
can then be instantiated to int or
string, permitting different instances of f
to be used in
differently-typed contexts.
It has long been known that let-generalization is unsound in the
presence of reference cells. For example, if the type inferred
for r
in the expression
let r = ref [] in r := [1]; "123"::!r This expression !r has type int list but is here used with type string listwere generalized from
'a list ref
to forall 'a. 'a list ref
, the
expression would have returned a list that contains a string and an
int, letting us apply string operations to an int, or vice versa. A
well-type program would have ``gone wrong.''
To ensure the type system soundness, there have been proposed
many various restrictions on let-generalization. The most widely implemented is
value restriction: the type inferred for a let-binding is
generalized only if the right-hand-side of the binding syntactically
has the form of a value. In other words, only values may have polymorphic
types. In our examples, the let-binding for f
(which
de-sugars to let f = fun () -> [] in ...
) binds f
to what syntactically
is a function, that is, a value. In contrast, the right-hand-side
of the binding to r
, ref []
, is syntactically a non-value
expression. Therefore, the type of r
is not generalized, prohibiting
the use of r
in differently-typed contexts, int list ref
vs string
list ref
. The type checker rejects our second example.
The value restriction has a clear intuition. We can type check our
first example without polymorphism, if we first inline the definition
of f
into its two use places. A polymorphic binding then may be
regarded as an `optimization', letting us type check f
once, where
it is defined, rather than at every place where f
is used. A polymorphic
type is an indication that the expression is inlineable; some compilers,
e.g., MLton, indeed inline all polymorphic expressions so that they can
be compiled without resorting to boxing. Effectful expressions (such
as ref []
) cannot be inlined while preserving dynamic behavior
as copying them replicates effects and hence is observable. Values are
inert and hence can be copied or shared at compiler's discretion,
without affecting dynamic behavior.
Alas, in a staged calculus with cross-stage persistence, value restriction or its variants (such as a relaxed value restriction) turn out insufficient to ensure type soundness. We demonstrate the unsoundness on a series of examples culminating in a segmentation fault. We develop our examples interactively, by submitting expressions to the top-level of a MetaOCaml interpreter of any recent version (309 or BER 002) and observing its responses. In the transcript below, the responses are indented.
The first example uses no staging and causes no controversy:
let f () = ref [] in f() := [1]; "123"::!(f()) - : string list = ["123"]The let-binding de-sugars into
let f = fun () -> ref [] in ...
whose right-hand side is syntactically a function. Value restriction
should allow generalization; both OCaml and MetaOCaml agree and
accept the expression. The indented line shows the result.What if we enclose the whole expression into MetaOCaml brackets? If an expression is well-typed, its code should be well-typed too. If we could manually enter an expression without type errors, we should be able to automatically generate that expression without type errors.
let c = .<let f () = ref [] in f() := [1]; "123"::!(f())>. val c : ('a, string list) code = .<let f_2 = fun () -> (ref ([])) in ((f_2 ()) := [1]); ("123" :: (! (f_2 ())))>. .! c;; - : string list = ["123"]MetaOCaml accepts the quoted expression from the first example. Running the code produces the result we have already seen.
Let us add an escape, or splice: .< let f () = .~(.<ref []>.) in ... >.
,
which reduces in one step to the second example.
let c = .<let f () = .~(.<ref []>.) in f() := [1]; "123"::!(f())>. val c : ('a, string list) code = .<let f_2 = fun () -> (ref ([])) in ((f_2 ()) := [1]); ("123" :: (! (f_2 ())))>. .! c;; - : string list = ["123"]Although the right-hand side of the
c
-binding is no longer a
(present-stage) value because of the splice, evaluating the right-hand-side
produces the code value of the second example, which is well-typed.
More formally, despite the splice, the second-stage binding to f
still
looks like a binding for a function. The (second) stage value
restriction should allow second-stage generalization. And it does, in
MetaOCaml. The result of running the code value c
is identical to that
of running the previous examples.We come to the final example, which looks as if it reduces to the earlier ones.
let c = .<let f () = .~(let x = ref [] in .<x>.) in f() := [1]; "123"::!(f())>. val c : ('a, string list) code = .<let f_2 = fun () -> (* cross-stage persistent value (as id: x) *) in ((f_2 ()) := [1]); ("123" :: (! (f_2 ())))>.The example is accepted by MetaOCaml; the code value
c
clearly
contains a cross-stage persistent value. We have created a reference cell at
the present stage and ``lifted'' the cell to the future-stage, letting
the generated code use the value as it is. We stress
that the cross-stage persistence is vital in practice:
without it, we have to write a staged version of
all standard library functions. The let-binding for f
is still
syntactically a binding for a function, so generalization seems justified.
If we try to run the generated code, we trip and fall:.! c;; segmentation faultAlthough the right-hand side of the future-stage
f
binding was
syntactically a functional value, fun () -> csprefval
, all applications of
that function return one and the same csprefval
. Therefore, sharing or
deep-copying of the function become observable -- the behavior that is
inconsistent with the inferred polymorphic type for the function. The
value restriction has a rarely mentioned premise: the only way to
produce reference cells is to evaluate an expression like ref
something
. There are no syntactic values of the type of reference
cells. Staging with cross-stage persistence can violate the premise:
reference cells are still created as the result of evaluating an
expression; that result, if lifted to the future stage, looks
syntactically like a value.The problem of restricting let-generalization to ensure soundness, considered closed long time ago, is thrown open in staged calculi.
Joint work with Chung-chieh Shan.
Jacques Garrigue: Relaxing the Value Restriction.
Proc. Int. Symposium on Functional and Logic Programming, Nara, April 2004. Springer-Verlag LNCS 2998, pp. 196--213. (extended version: RIMS Preprint 1444)
< http://www.math.nagoya-u.ac.jp/~garrigue/papers/morepoly-long.pdf >
The paper gives a good survey of approaches to restrict let-generalization to ensure its soundness.
run
and cross-stage persistencelambda^{a}_{v1}
is the model of MetaOCaml restricted to two levels. This is
a call-by-value two-level lambda-calculus that
supports the manipulation and the splicing-in of the open code,
the running of the closed code, and cross-staged persistence. The calculus
implements Taha and Nielsen's `environment classifiers' to prevent
attempts to run open code. The calculus is a small variation of
the one presented in Taha and Nielsen's POPL2003 paper.The calculus -- its syntax, dynamic and static semantics -- is implemented in Twelf. We can enter terms, infer their types and see their values. The Twelf code includes many sample terms and the examples of type inference and evaluation. The implementation was used to write all examples in the ``Closing the Stage'' (PEPM 2008) paper: The calculus was the source language of the translation described in the paper.
In addition to implementing the calculus, the Twelf code proves several meta-theoretical properties: any non-value term can be decomposed into a possibly open pre-value and the possibly binding context; primitive reductions preserve types. The proofs are very challenging: splices let evaluation happen under a future-stage lambda. The evaluation context therefore can cross the arbitrary number of dependent binders: variable binders include classifiers that must be bound, too, at that point.
lambda-am1.elf [34K]
Twelf code
The code includes many examples of type-checking and evaluating staged terms
util.elf [<1K]
Common utilities: Natural numbers and their addition and equality
Here is the simplest example of scope extrusion in MetaOCaml caused by the effect of mutating a state:
# let code = let x = ref .<1>. in let _ = .<fun v -> .~(x := .<v>.; .<()>.)>. in !x;; val code: ('a, int) code = .<v_1>. # .!code;; Unbound value v_1 Exception: Trx.TypeCheckingError.We have managed to build a piece of code with literally a free variable. Evaluating this code (by MetaOCaml's `run' operation
.!
) causes a paradoxical type error at run-time. The type-checker has accepted the code
that it should not have. We have seen such scope extrusion arising from honest mistakes of let-insertion in real program generation with effects. No staged language today can statically prevent such mistakes.Our goal is to make program generation convenient and safe, in particular, to statically prevent errors like scope extrusion. Developing a sound type system of staged code with effects requires a suitable calculus. To model real staged programming languages like MetaOCaml, we need a call by value calculus that supports splicing of open code, running the closed code, and cross-stage persistence. Modeling effects, especially control effects such as delimited continuations, is better with small-step operational semantics. Many formal calculi for staged programming have evolved over the last couple of decades. Here is a sample:
Our goal for this project is to make, with the benefit of hindsight, the existing staged calculi more uniform and their features more orthogonal to each other, so that they are easier to study, mechanize, and extend (for example, to add side effects). To be more precise, we aim to:
Joint work with Chung-chieh Shan and Yukiyoshi Kameyama.
Two-level staged calculus with environment classifiers, run
and cross-stage persistence
This calculus is the result of our quest. The calculus is a close model of MetaOCaml.
Shifting the Stage: Staging with Delimited Control
A simpler calculus, without cross-stage persistence and first-class run
. The calculus is significantly easier to work with. It let us add delimited control to staging. The meta-theory of the resulting calculus has been fully mechanized, with mechanically checked proofs of progress and preservation.
Our solution represents contexts outside-in and uses dependent types to describe the binding structure of contexts and the corresponding structure of open terms. We convinced Twelf that our our functions to decompose a term into an (open) redex and its context, and to plug an open term into its closing context are total. This totality suggests that our types adequately represent ordered dependent sequences of bindings, be they needed by an expression or provided by a context. These focusing and zipping functions let us specify the first small-step semantics for staging.
The challenge remains to represent contexts inside-out while expressing its binding structure, in particular how the continuation of a staged evaluator may ``bind off'' a later-stage variable.
Joint work with Chung-chieh Shan.
oleg-at-okmij.org
Your comments, problem reports, questions are very welcome!
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