A Verified Compiler from Isabelle/HOL
to CakeML
Lars Hupel
and Tobias Nipkow
Technische Universit¨at M¨unchen, Munich, Germany
lars.hupel@tum.de, nipkow@in.tum.de
Abstract. Many theorem provers can generate functional programs
from definitions or proofs. However, this code generation needs to be
trusted. Except for the HOL4 system, which has a proof producing
code generator for a subset of ML. We go one step further and provide
a verified compiler from Isabelle/HOL to CakeML. More precisely we
combine a simple proof producing translation of recursion equations in
Isabelle/HOL into a deeply embedded term language with a fully verified
compilation chain to the target language CakeML.
Keywords: Isabelle
· CakeML · Compiler
Higher-order term rewriting
1 Introduction
Many theorem provers have the ability to generate executable code in some (typ-
ically functional) programming language from definitions, lemmas and proofs
(e.g. [6,8,9,12,16,27,37]). This makes code generation part of the trusted kernel
of the system. Myreen and Owens [30] closed this gap for the HOL4 system: they
have implemented a tool that translates from HOL4 into CakeML, a subset of
SML, and proves a theorem stating that a result produced by the CakeML code
is correct w.r.t. the HOL functions. They also have a verified implementation of
CakeML [24,40]. We go one step further and provide a once-and-for-all verified
compiler from (deeply embedded) function definitions in Isabelle/HOL [32,33]
into CakeML proving partial correctness of the generated CakeML code w.r.t.
the original functions. This is like the step from dynamic to static type checking.
It also means that preconditions on the input to the compiler are explicitly given
in the correctness theorem rather than implicitly by a failing translation. To the
best of our knowledge this is the first verified (as opposed to certifying) compiler
from function definitions in a logic into a programming language.
Our compiler is composed of multiple phases and in principle applicable to
other languages than Isabelle/HOL or even HOL:
The Author(s) 2018
A. Ahmed (Ed.): ESOP 2018, LNCS 10801, pp. 999–1026, 2018.
1000 L. Hupel and T. Nipkow
We erase types right away. Hence the type system of the source language is
We merely assume that the source language has a semantics based on equa-
tional logic.
The compiler operates in three stages:
1. The preprocessing phase eliminates features that are not supported by our
compiler. Most importantly, dictionary construction eliminates occurrences
of type classes in HOL terms. It introduces dictionary datatypes and new
constants and proves the equivalence of old and new constants (Sect. 7).
2. The deep embedding lifts HOL terms into terms of type term, a HOL model
of HOL terms. For each constant c (of arbitrary type) it defines a constant c
of type term and proves a theorem that expresses equivalence (Sect. 3).
3. There are multiple compiler phases that eliminate certain constructs from
the term type, until we arrive at the CakeML expression type. Most phases
target a different intermediate term type (Sect. 5).
The first two stages are preprocessing, are implemented in ML and produce
certificate theorems. Only these stages are specific to Isabelle. The third (and
main) stage is implemented completely in the logic HOL, without recourse to
ML. Its correctness is verified once and for all.
2 Related Work
There is existing work in the Coq [2,15]andHOL[30] communities for proof
producing or verified extraction of functions defined in the logic. Anand et al. [2]
present work in progress on a verified compiler from Gallina (Coq’s specification
language) via untyped intermediate languages to CompCert C light. They plan
to connect their extraction routine to the CompCert compiler [26].
Translation of type classes into dictionaries is an important feature of Haskell
compilers. In the setting of Isabelle/HOL, this has been described by Wenzel
[44] and Krauss et al. [23]. Haftmann and Nipkow [17] use this construction to
compile HOL definitions into target languages that do not support type classes,
e.g. Standard ML and OCaml. In this work, we provide a certifying translation
that eliminates type classes inside the logic.
Compilation of pattern matching is well understood in literature [3,36,38].
In this work, we contribute a transformation of sets of equations with pattern
matching on the left-hand side into a single equation with nested pattern match-
ing on the right-hand side. This is implemented and verified inside Isabelle.
Besides CakeML, there are many projects for verified compilers for functional
programming languages of various degrees of sophistication and realism (e.g.
All Isabelle definitions and proofs can be found on the paper website: https://
A Verified Compiler from Isabelle/HOL to CakeML 1001
[4,11,14]). Particularly modular is the work by Neis et al. [31] on a verified
compiler for an ML-like imperative source language. The main distinguishing
feature of our work is that we start from a set of higher-order recursion equations
with pattern matching on the left-hand side rather than a lambda calculus with
pattern matching on the right-hand side. On the other hand we stand on the
shoulders of CakeML which allows us to bypass all complications of machine
code generation. Note that much of our compiler is not specific to CakeML and
that it would be possible to retarget it to, for example, Pilsner abstract syntax
with moderate effort.
Finally, Fallenstein and Kumar [13] have presented a model of HOL inside
HOL using large cardinals, including a reflection proof principle.
3 Deep Embedding
Starting with a HOL definition, we derive a new, reied definition in a deeply
embedded term language depicted in Fig. 1a. This term language corresponds
closely to the term datatype of Isabelle’s implementation (using de Bruijn indices
[10]), but without types and schematic variables.
To establish a formal connection between the original and the reified defini-
tions, we use a logical relation, a concept that is well-understood in literature
[20] and can be nicely implemented in Isabelle using type classes. Note that the
use of type classes here is restricted to correctness proofs; it is not required for
the execution of the compiler itself. That way, there is no contradiction to the
elimination of type classes occurring in a previous stage.
Notation. We abbreviate App tuto t $ u and Abs t to Λt. Other term types
introduced later in this paper use the same conventions. We reserve λ for abstrac-
tions in HOL itself. Typing judgments are written with a double colon: t :: τ.
Embedding Operation. Embedding is implemented in ML. We denote this oper-
ation using angle brackets: t, where t is an arbitrary HOL expression and the
result t is a HOL value of type term. It is a purely syntactic transformation,
without preliminary evaluation or reduction, and it discards type information.
The following examples illustrate this operation and typographical conventions
concerning variables and constants:
x = Free "x" f = Const "f" λx. f x = Λ (f $ Bound 0)
Small-Step Semantics. Figure 1b specifies the small-step semantics for term.Itis
reminiscent of higher-order term rewriting, and modelled closely after equality in
HOL. The basic idea is that if the proposition t = u
can be proved equationally
in HOL (without symmetry), then R t−
u holds (where R :: (term ×
term) set ). We call R the rule set. It is the result of translating a set of defining
equations lhs = rhs into pairs (lhs , rhs) R.
1002 L. Hupel and T. Nipkow
datatype term =
Const string |
Free string |
Abs term |
Bound nat |
App term term
(a) Abstract syntax of
de Bruijn terms
(lhs, rhs) R match lhs t = Some σ
R t −→ subst σ rhs
closed t
R (Λt)$t
−→ t[t
R t −→ t
R t $ u −→ t
$ u
R u −→ u
R t $ u −→ t $ u
(b) Small-step semantics
Fig. 1. Basic syntax and semantics of the term type
Rule Step performs a rewrite step by picking a rewrite rule from R and
rewriting the term at the root. For that purpose, match and subst are (mostly)
standard first-order matching and substitution (see Sect. 4 for details).
Rule Beta performs β-reduction. Type term represents bound variables by
de Bruijn indices. The notation t[t
] represents the substitution of the outermost
bound variable in t with t
Our semantics does not constitute a fully-general higher-order term rewrit-
ing system, because we do not allow substitution under binders. For de Bruijn
terms, this would pose no problem, but as soon as we introduce named bound
variables, substitution under binders requires dealing with capture. To avoid this
altogether, all our semantics expect terms that are substituted into abstractions
to be closed. However, this does not mean that we restrict ourselves to any par-
ticular evaluation order. Both call-by-value and call-by-name can be used in the
small-step semantics. But later on, the target semantics will only use call-by-
Embedding Relation. We denote the concept that an embedded term t corre-
sponds to a HOL term a of type τ w.r.t. rule set R with the syntax R t a.
If we want to be explicit about the type, we index the relation:
For ground types, this can be defined easily. For example, the following two
rules define
R 0≈
R t≈
R Suc t≈
Suc n
Definitions of for arbitrary datatypes without nested recursion can be derived
mechanically in the same fashion as for nat, where they constitute one-to-
one relations. Note that for ground types, ignores R. The reason why is
parametrized on R will become clear in a moment.
For function types, we follow Myreen and Owen’s approach [30]. The state-
ment R t f can be interpreted as t $ a can be rewritten to fa for
all a”. Because this might involve applying a function definition from R,the
relation must be indexed by the rule set. As a notational convenience, we define
A Verified Compiler from Isabelle/HOL to CakeML 1003
another relation R t x to mean that there is a t
such that R t −→
R t
x. Using this notation, we formally define for functions as follows:
R t f (ux.R u x R t $ u fx)
Example. As a running example, we will use the map function on lists:
map f [] = []
map f (x # xs)=fx# map fxs
The result of embedding this function is a set of rules map
map’ =
{(Const ”List.list.map” $ Free ”f” $ (Const ”List.list.Cons” $ Free ”x21” $ Free ”x22”),
Const ”List.list.Cons” $ (Free ”f” $ Free ”x21”) $ ...),
(Const ”List.list.map” $ Free ”f” $ Const ”List.list.Nil”,
Const ”List.list.Nil”)}
together with the theorem map
Const "List.list.map" map,whichis
proven by simple induction over map. Constant names like "List.list.map"
come from the fully-qualified internal names in HOL.
The induction principle for the proof arises from the use of the fun command
that is used to define recursive functions in HOL [22]. But the user is also allowed
to specify custom equations for functions, in which case we will use heuristics
to generate and prove the appropriate induction theorem. For simplicity, we
will use the term (defining) equation uniformly to refer to any set of equations,
either default ones or ones specified by the user. Embedding partially-specified
functions in particular, proving the certificate theorem about them is cur-
rently not supported. In the future, we plan to leverage the domain predicate as
produced by fun to generate conditional theorems.
4 Terms, Matching and Substitution
The compiler transforms the initial term type (Fig. 1a) through various inter-
mediate stages. This section gives an overview and introduces necessary
Preliminaries. The function arrow in HOL is . The cons operator on lists is
the infix #.
Throughout the paper, the concept of mappings is pervasive: We use the
type notation αβto denote a function α β option. In certain contexts,
a mapping may also be called an environment. We write mapping literals using
brackets: [a x, b y,...]. If it is clear from the context that σ is defined on
a, we often treat the lookup σaas returning an x :: β.
The functions dom :: (αβ) α set and range :: (αβ) β set return
the domain and range of a mapping, respectively.
1004 L. Hupel and T. Nipkow
Dropping entries from a mapping is denoted by σ k, where σ is a mapping
and k is either a single key or a set of keys. We use σ
σ to denote that σ
a sub-mapping of σ,thatis,dom σ
dom σ and a dom σ
a = σa.
Merging two mappings σ and ρ is denoted with σ ++ ρ. It constructs a new
mapping with the union domain of σ and ρ.Entriesfromρ override entries from
σ.Thatis,ρ σ ++ ρ holds, but not necessarily σ σ ++ ρ.
All mappings and sets are assumed to be finite. In the formalization, this is
enforced by using subtypes of and set. Note that one cannot define datatypes
by recursion through sets for cardinality reasons. However, for finite sets, it
is possible. This is required to construct the various term types. We leverage
facilities of Blanchette et al.’s datatype command to define these subtypes [7].
Standard Functions. All type constructors that we use (, set, list, option, ...)
support the standard operations map and rel. For lists, map is the regular covariant
map. For mappings, the function has the type (β γ) (αβ) (αγ).
It leaves the domain unchanged, but applies a function to the range of the
Function rel
lifts a binary predicate P :: α α bool to the type construc-
tor τ. We call this lifted relation the relator for a particular type.
For datatypes, its definition is structural, for example:
P [] []
P xs ys Pxy
P (x # xs)(y # ys)
For sets and mappings, the definition is a little bit more subtle.
Definition 1 (Set relator). For each element a A, there must be a corre-
sponding element b B such that Pab, and vice versa. Formally:
PAB (x A. y B. P x y) (y B. x A. P x y)
Definition 2 (Mapping relator). For each a, maand namust be related
according to rel
P . Formally:
Pmn (a. rel
P (ma)(na))
Term Types. There are four distinct term types: term, nterm, pterm,andsterm.
All of them support the notions of free variables, matching and substitution. Free
variables are always a finite set of strings. Matching a term against a pattern
yields an optional mapping of type string αfrom free variable names to terms.
Note that the type of patterns is itself term instead of a dedicated pattern
type. The reason is that we have to subject patterns to a linearity constraint
anyway and may use this constraint to carve out the relevant subset of terms:
Definition 3. Atermislinear if there is at most one occurrence of any variable,
it contains no abstractions, and in an application f $ x, f must not be a free
variable. The HOL predicate is called linear :: term bool.
A Verified Compiler from Isabelle/HOL to CakeML 1005
Because of the similarity of operations across the term types, they are all
instances of the term type class. Note that in Isabelle, classes and types live
in different namespaces. The term type and the term type class are separate
Definition 4. A term type τ supports the operations match :: term τ
(string τ), subst :: (string τ) τ τ and frees :: τ string set.We
also define the following derived functions:
matchs matches a list of patterns and terms sequentially, producing a single
closed t is an abbreviation for frees t =
closed σ is an overloading of closed, denoting that all values in a mapping are
Additionally, some (obvious) axioms have to be satisfied. We do not strive to
fully specify an abstract term algebra. Instead, the axioms are chosen according
to the needs of this formalization.
A notable deviation from matching as discussed in term rewriting literature
is that the result of matching is only well-defined if the pattern is linear.
Definition 5. An equation is a pair of a pattern (left-hand side) andaterm
(right-hand side). The pattern is of the form f $p
,wheref is a constant
(i.e. of the form Const name). We refer to both f or name interchangeably as
the function symbol of the equation.
Following term rewriting terminology, we sometimes refer to an equation as rule.
4.1 De Bruijn terms (term)
The definition of term is almost an exact copy of Isabelle’s internal term type,
with the notable omissions of type information and schematic variables (Fig. 1a).
The implementation of β-reduction is straightforward via index shifting of bound
4.2 Named Bound Variables (nterm)
datatype nterm = Nconst string | Nvar string | Nabs string nterm | Napp nterm nterm
The nterm type is similar to term, but removes the distinction between bound
and free variables. Instead, there are only named variables. As mentioned in the
previous section, we forbid substitution of terms that are not closed in order
to avoid capture. This is also reflected in the syntactic side conditions of the
correctness proofs (Sect. 5.1).
1006 L. Hupel and T. Nipkow
4.3 Explicit Pattern Matching (pterm)
datatype pterm =
Pconst string | Pvar string | Pabs ((term × pterm) set) | Papp pterm pterm
Functions in HOL are usually defined using implicit pattern matching, that is,
the terms p
occurring on the left-hand side f p
... p
of an equation must
be constructor patterns. This is also common among functional programming
languages like Haskell or OCaml. CakeML only supports explicit pattern match-
ing using case expressions. A function definition consisting of multiple defining
equations must hence be translated to the form f = λx. case x of ....The
elimination proceeds by iteratively removing the last parameter in the block of
equations until none are left.
In our formalization, we opted to combine the notion of abstraction and case
expression, yielding case abstractions, represented as the Pabs constructor. This
is similar to the fn construct in Standard ML, which denotes an anonymous
function that immediately matches on its argument [28]. The same construct
also exists in Haskell with the LambdaCase language extension. We chose this
representation mainly for two reasons: First, it allows for a simpler language
grammar because there is only one (shared) constructor for abstraction and case
expression. Second, the elimination procedure outlined above does not have to
introduce fresh names in the process. Later, when translating to CakeML syntax,
fresh names are introduced and proved correct in a separate step.
The set of pairs of pattern and right-hand side inside a case abstraction is
referred to as clauses. As a short-hand notation, we use Λ{p
4.4 Sequential Clauses (sterm)
datatype sterm =
Sconst string | Svar string | Sabs ((term × sterm) list) | Sapp sterm sterm
In the term rewriting fragment of HOL, the order of rules is not significant. If a
rule matches, it can be applied, regardless when it was defined or proven. This
is reflected by the use of sets in the rule and term types. For CakeML, the rules
need to be applied in a deterministic order, i.e. sequentially. The sterm type only
differs from pterm by using list instead of set. Hence, case abstractions use list
brackets: Λ[p
4.5 Irreducible Terms (value)
CakeML distinguishes between expressions and values. Whereas expressions may
contain free variables or β-redexes, values are closed and fully evaluated. Both
have a notion of abstraction, but values differ from expressions in that they
contain an environment binding free variables.
Consider the expression (λx.λy.x)(λz.z), which is rewritten (by β-reduction)
to λy.λz.z. Note how the bound variable x disappears, since it is replaced. This
A Verified Compiler from Isabelle/HOL to CakeML 1007
is contrary to how programming languages are usually implemented: evaluation
does not happen by substituting the argument term t for the bound variable
x, but by recording the binding x → t in an environment [24]. A pair of an
abstraction and an environment is usually called a closure [25,41].
In CakeML, this means that evaluation of the above expression results in the
(λy.x, ["x" → (λz.z, [])])
Note the nested structure of the closure, whose environment itself contains a
To reflect this in our formalization, we introduce a type value of values (expla-
nation inline):
datatype value =
(constructor value: a data constructor applied to multiple values )
Vconstr string (value list) |
(closure: clauses combined with an environment mapping variables to values )
Vabs ((term × sterm) list) (string value) |
(recursive closures: a group of mutually recursive function bodies with an environment )
Vrecabs (string ((term × sterm) list)) string (string value)
The above example evaluates to the closure:
"x" → Vabs [z⇒z][]
The third case for recursive closures only becomes relevant when we conflate
variables and constants. As long as the rule set rs is kept separate, recursive calls
are straightforward: the appropriate definition for the constant can be looked up
there. CakeML knows no such distinction between constants and variables, hence
everything has to reside in a single environment σ.
Consider this example of odd and even:
odd 0=False even 0=True
odd (Suc n)=even n even (Suc n)=odd n
When evaluating the term odd k, the definitions of even and odd themselves
must be available in the environment captured in the definition of odd. However,
it would be cumbersome in HOL to construct such a Vabs that refers to itself.
Instead, we capture the expressions used to define odd and even in a recursive
closure. Other encodings might be possible, but since we are targeting CakeML,
we are opting to model it in a similar way as its authors do.
For the above example, this would result in the following global environment:
["odd" → Vrecabs css "odd" [], "even" → Vrecabs css "even" []]
where css =["odd" → [0⇒False ,
Suc n⇒even n],
"even" → [0⇒True , Suc n⇒odd n]]
1008 L. Hupel and T. Nipkow
Note that in the first line, the right-hand sides are values, but in css, they
are expressions. The additional string argument of Vrecabs denotes the selected
function. When evaluating an application of a recursive closure to an argument
(β-reduction), the semantics adds all constituent functions of the closure to the
environment used for recursive evaluation.
5 Intermediate Semantics and Compiler Phases
In this section, we will discuss the progression from de Bruijn based term lan-
guage with its small-step semantics given in Fig. 1a to the final CakeML seman-
tics. The compiler starts out with terms of type term and applies multiple
phases to eliminate features that are not present in the CakeML source language.
de Bruijn
Named bound
Explicit pattern
constructors :: string set (shared by all phases)
R :: (term × term) set, t, t
:: term
R t −→ t
(Figure 1b)
R :: (term × nterm) set, t, t
:: nterm
R t −→ t
(Figure 3)
R :: (string × pterm) set, t, t
:: pterm
R t −→ t
(Figure 4)
rs :: (string × sterm) list, t, t
:: sterm
rs t −→ t
(Figure 5)
rs :: (string × sterm) list, σ :: string sterm
t, u :: sterm
rs t u (Figure 6)
rs :: (string × value) list, σ :: string value
t :: sterm, u :: value
rs t u (Figure 7)
σ :: string value
t :: sterm, u :: value
σ t u (Figure 8)
Phase/Refinement Types & Semantics
Theorem 1
see §5.3
see §5.4
Theorem 2
Theorem 1
Theorem 4
compiler phase; semantics refinement
semantics belonging to the phase; semantics relation
Fig. 2. Intermediate semantics and compiler phases
A Verified Compiler from Isabelle/HOL to CakeML 1009
Types term, nterm and pterm each have a small-step semantics only. Type sterm
has a small-step and several intermediate big-step semantics that bridge the gap
to CakeML. An overview of the intermediate semantics and compiler phases is
depicted in Fig. 2. The left-hand column gives an overview of the different phases.
The right-hand column gives the types of the rule set and the semantics for each
phase; you may want to skip it upon first reading.
(lhs, rhs) R match lhs t = Some σ
R t −→ subst σ rhs
closed t
R (Λx. t)$t
−→ subst
x → t
Fig. 3. Small-step semantics for nterm with named bound variables
5.1 Side Conditions
All of the following semantics require some side conditions on the rule set. These
conditions are purely syntactic. As an example we list the conditions for the
correctness of the first compiler phase:
Patterns must be linear, and constructors in patterns must be fully applied.
Definitions must have at least one parameter on the left-hand side (Sect. 5.6).
The right-hand side of an equation refers only to free variables occurring in
patterns on the left-hand side and contain no dangling de Bruijn indices.
There are no two defining equations lhs = rhs
and lhs = rhs
such that
= rhs
For each pair of equations that define the same constant, their arity must be
equal and their patterns must be compatible (Sect. 5.3).
There is at least one equation.
Variable names occurring in patterns must not overlap with constant names
(Sect. 5.7).
Any occurring constants must either be defined by an equation or be a con-
The conditions for the subsequent phases are sufficiently similar that we do not
list them again.
In the formalization, we use named contexts to fix the rules and assump-
tions on them (locales in Isabelle terminology). Each phase has its own locale,
together with a proof that after compilation, the preconditions of the next phase
are satisfied. Correctness proofs assume the above conditions on R and similar
conditions on the term that is reduced. For brevity, this is usually omitted in
our presentation.
1010 L. Hupel and T. Nipkow
5.2 Naming Bound Variables: From term to nterm
Isabelle uses de Bruijn indices in the term language for the following two rea-
sons: For substitution, there is no need to rename bound variables. Additionally,
α-equivalent terms are equal. In implementations of programming languages,
these advantages are not required: Typically, substitutions do not happen inside
abstractions, and there is no notion of equality of functions. Therefore CakeML
uses named variables and in this compilation step, we get rid of de Bruijn indices.
The “named” semantics is based on the nterm type. The rules that are
changed from the original semantics (Fig. 1b) are given in Fig. 3 (Fun and Arg
remain unchanged). Notably, β-reduction reuses the substitution function.
For the correctness proof, we need to establish a correspondence between
termsandnterms. Translation from nterm to term is trivial: Replace bound
variables by the number of abstractions between occurrence and where they
were bound in, and keep free variables as they are. This function is called
to term.
The other direction is not unique and requires introduction of fresh names
for bound variables. In our formalization, we have chosen to use a monad to
produce these names. This function is called term
to nterm. We can also prove
the obvious property nterm
to term (term to nterm t)=t, where t is a term
without dangling de Bruijn indices.
Generation of fresh names in general can be thought of as picking a string
that is not an element of a (finite) set of already existing names. For Isabelle,
the Nominal framework [42,43] provides support for reasoning over fresh names,
but unfortunately, its definitions are not executable.
Instead, we chose to model generation of fresh names as a monad α fresh
with the following primitive operations in addition to the monad operations:
run:: α fresh string set α
name:: string fresh
In our implementation, we have chosen to represent α fresh as roughly isomorphic
to the state monad.
Compilation of a rule set proceeds by translation of the right-hand side of all
compile R = {(p, term
to nterm t) | (p, t) R}
The left-hand side is left unchanged for two reasons: function match expects an
argument of type term (see Sect. 4), and patterns do not contain abstractions or
bound variables.
Theorem 1 (Correctness of compilation). Assuming a step can be taken
with the compiled rule set, it can be reproduced with the original rule set.
compile R t −→ u closed t
R nterm to term t −→ nterm to term u
We prove this by induction over the semantics (Fig. 3).
A Verified Compiler from Isabelle/HOL to CakeML 1011
(pat , rhs) C match pat t = Some σ closed t
R (Λ C)$t subst σ rhs
(name, rhs) R
R Pconst name rhs
Fig. 4. Small-step semantics for pterm with pattern matching
5.3 Explicit Pattern Matching: From nterm to pterm
Usually, functions in HOL are defined using implicit pattern matching, that is,
the left-hand side of an equation is of the form f p
... p
, where the p
patterns over datatype constructors. For any given function f, there may be
multiple such equations. In this compilation step, we transform sets of equations
for f defined using implicit pattern matching into a single equation for f of the
form f = Λ C, where C is a set of clauses.
The strategy we employ currently requires successive elimination of a single
parameter from right to left, in a similar fashion as Slind’s pattern matching
compiler [38, Sect. 3.3.1]. Recall our running example (map). It has arity 2. We
omit the brackets  for brevity. First, the list parameter gets eliminated:
map f = λ [] []
| x # xs fx# map fxs
Finally, the function parameter gets eliminated:
map = λf
λ [] []
| x # xs fx# map fxs
This has now arity 0 and is defined by a twice-nested abstraction.
Semantics. The target semantics is given in Fig. 4 (the Fun and Arg rules
from previous semantics remain unchanged). We start out with a rule set R that
allows only implicit pattern matching. After elimination, only explicit pattern
matching remains. The modified Step rule merely replaces a constant by its
definition, without taking arguments into account.
Restrictions. For the transformation to work, we need a strong assumption
about the structure of the patterns p
to avoid the following situation:
map f [] = []
map g (x # xs)=gx# map gxs
Through elimination, this would turn into:
map = λf
λ [] []
| g
λx# xs fx# map fxs
1012 L. Hupel and T. Nipkow
(name, rhs) R
R Sconst name rhs
match cs t = Some (σ, rhs) closed t
R (Λ cs)$t subst σ rhs
Fig. 5. Small-step semantics for sterm
Even though the original equations were non-overlapping, we suddenly
obtained an abstraction with two overlapping patterns. Slind observed a similar
problem [38, Sect. 3.3.2] in his algorithm. Therefore, he only permits uniform
equations, as defined by Wadler [36, Sect. 5.5]. Here, we can give a formal char-
acterization of our requirements as a computable function on pairs of patterns:
fun pat compat :: term term bool where
compat (t
$ t
$ u
) pat compat t
= u
pat compat t
compat tu(overlapping tut = u)
This compatibility constraint ensures that any two overlapping patterns (of the
same column) p
and p
are equal and are thus appropriately grouped together
in the elimination procedure. We require all defining equations of a constant to be
mutually compatible. Equations violating this constraint will be flagged during
embedding (Sect. 3), whereas the pattern elimination algorithm always succeeds.
While this rules out some theoretically possible pattern combinations (e.g.
the diagonal function [36, Sect. 5.5]), in practice, we have not found this to be a
problem: All of the function definitions we have tried (Sect. 8) satisfied pattern
compatibility (after automatic renaming of pattern variables). As a last resort,
the user can manually instantiate function equations. Although this will always
lead to a pattern compatible definition, it is not done automatically, due to the
potential blow-up.
Discussion. Because this compilation phase is both non-trivial and has some
minor restrictions on the set of function definitions that can be processed, we
may provide an alternative implementation in the future. Instead of eliminat-
ing patterns from right to left, patterns may be grouped in tuples. The above
example would be translated into:
map = λ (f, []) []
| (f, x # xs) fx# map fxs
We would then leave the compilation of patterns for the CakeML compiler, which
has no pattern compatibility restriction.
The obvious disadvantage however is that this would require the knowledge
of a tuple type in the term language which is otherwise unaware of concrete
5.4 Sequentialization: From pterm to sterm
The semantics of pterm and sterm differ only in rule Step and Beta. Figure 5
shows the modified rules. Instead of any matching clause, the first matching
clause in a case abstraction is picked.
A Verified Compiler from Isabelle/HOL to CakeML 1013
For the correctness proof, the order of clauses does not matter: we only need
to prove that a step taken in the sequential semantics can be reproduced in the
unordered semantics. As long as no rules are dropped, this is trivially true. For
that reason, the compiler orders the clauses lexicographically. At the same time
the rules are also converted from type (string × pterm) set to (string × sterm) list.
Below, rs will always denote a list of the latter type.
(name, rhs) rs
rs Sconst name rhs
Va r
σ name = Some v
rs Svar name v
rs Λ cs Λ [(pat , subst (σ frees pat ) t | (pat ,t) cs]
rs t Λ cs
rs u u
first match cs u
= Some (σ
, rhs) rs++ σ
rhs v
rs t $ u v
name constructors rs t
··· rs t
rs Sconst name $ t
$ ...$ t
Sconst name $ u
$ ...$ u
Fig. 6. Big-step semantics for sterm
5.5 Big-Step Semantics for sterm
This big-step semantics for sterm is not a compiler phase but moves towards
the desired evaluation semantics. In this first step, we reuse the sterm type for
evaluation results, instead of evaluating to the separate type value. This allows
us to ignore environment capture in closures for now.
All previous −→ relations were parametrized by a rule set. Now the big-step
predicate is of the form rs t t
where σ :: string sterm is a variable
This semantics also introduces the distinction between constructors and
defined constants. If C is a constructor, the term C t
... t
is evaluated to
C t
... t
where the t
are the results of evaluating the t
The full set of rules is shown in Fig. 6. They deserve a short explanation:
Const. Constants are retrieved from the rule set rs.
Var. Variables are retrieved from the environment σ.
Abs. In order to achieve the intended invariant, abstractions are evaluated to
their fully substituted form.
Comb. Function application t $ u first requires evaluation of t into an abstrac-
tion Λ cs and evaluation of u into an arbitrary term u
. Afterwards, we look
for a clause matching u
in cs, which produces a local variable environment
, possibly overwriting existing variables in σ. Finally, we evaluate the right-
hand side of the clause with the combined global and local variable environ-
Constr. For a constructor application C t
..., evaluate all t
. The set con-
structors is an implicit parameter of the semantics.
1014 L. Hupel and T. Nipkow
(name, rhs) rs
rs Sconst name rhs
Va r
σ name = Some v
rs Svar name v
rs Λ cs Vabs cs σ
rs t Vabs cs σ
rs u v first match cs v = Some (σ
, rhs) rs
++ σ
rhs v
rs t $ u v
rs t Vrecabs css name σ
css name = Some cs rs u v
match cs v = Some (σ
, rhs) rs
++ σ
rhs v
rs t $ u v
name constructors rs t
··· rs t
rs Sconst name $ t
$ ...$ t
Vconstr name [v
Fig. 7. Evaluation semantics from sterm to value
Lemma 1 (Closedness invariant). If σ contains only closed terms, frees t
dom σ and rs t t
, then t
is closed.
Correctness of the big-step w.r.t. the small-step semantics is proved easily by
induction on the former:
Lemma 2. For any closed environment σ satisfying frees t dom σ,
rs t u rs subst σt−→
By setting σ = [], we obtain:
Theorem 2 (Correctness). rs, [] t u closed t rs t −→
5.6 Evaluation Semantics: Refining sterm to value
At this point, we introduce the concept of values into the semantics, while still
keeping the rule set (for constants) and the environment (for variables) separate.
The evaluation rules are specified in Fig. 7 and represent a departure from the
original rewriting semantics: a term does not evaluate to another term but to an
object of a different type, a value. We still use as notation, because big-step
and evaluation semantics can be disambiguated by their types.
The evaluation model itself is fairly straightforward. As explained in Sect. 4.5,
abstraction terms are evaluated to closures capturing the current variable envi-
ronment. Note that at this point, recursive closures are not treated differently
from non-recursive closures. In a later stage, when rs and σ are merged, this
distinction becomes relevant.
A Verified Compiler from Isabelle/HOL to CakeML 1015
We will now explain each rule that has changed from the previous semantics:
Abs. Abstraction terms are evaluated to a closure capturing the current
Comb. As before, in an application t$u, t must evaluate to a closure Vabs cs σ
The evaluation result of u is then matched against the clauses cs, producing
an environment σ
. The right-hand side of the clause is then evaluated using
++ σ
; the original environment σ is effectively discarded.
RecComb. Similar as above. Finding the matching clause is a two-step process:
First, the appropriate clause list is selected by name of the currently active
function. Then, matching is performed.
Constr. As before, for an n-ary application C t
..., where C is a data con-
structor, we evaluate all t
. The result is a Vconstr value.
Conversion Between sterm and value. To establish a correspondence between
evaluating a term to an sterm and to a value, we apply the same trick as in
Sect. 5.2. Instead of specifying a complicated relation, we translate value back
to sterm: simply apply the substitutions in the captured environments to the
The translation rules for Vabs and Vrecabs are kept similar to the Abs rule
from the big-step semantics (Fig. 6). Roughly speaking, the big-step semantics
always keeps terms fully substituted, whereas the evaluation semantics defers
Similarly to Sect. 5.2, we can also define a function sterm
to value :: sterm
value and prove that one function is the inverse of the other.
Matching. The value type, instead of using binary function application as all
other term types, uses n-ary constructor application. This introduces a concep-
tual mismatch between (binary) patterns and values. To make the proofs easier,
we introduce an intermediate type of n-ary patterns. This intermediate type can
be optimized away by fusion.
Correctness. The correctness proof requires a number of interesting lemmas.
Lemma 3 (Substitution before evaluation). Assuming that a term t can
be evaluated to a value u given a closed environment σ, it can be evaluated to
the same value after substitution with a sub-environment σ
. Formally: rs
t u σ
σ rs subst σ
t u
This justifies the “pre-substitution” exhibited by the Abs rule in the big-step
semantics in contrast to the environment-capturing Abs rule in the evaluation
Theorem 3 ( Correctness). Let σ beaclosedenvironmentandt a term which
only contains free variables in dom σ. Then, an evaluation to a value rs t v
can be reproduced in the big-step semantics as rs
, map value to sterm σ t
to sterm v, where rs
=[(name, value to sterm rhs) | (name, rhs) rs].
1016 L. Hupel and T. Nipkow
Instantiating the Correctness Theorem. The correctness theorem states
that, for any given evaluation of a term t with a given environment rs con-
taining values, we can reproduce that evaluation in the big-step semantics using
a derived list of rules rs
and an environment σ
containing sterms that are gen-
erated by the value
to sterm function. But recall the diagram in Fig. 2.Inour
scenario, we start with a given rule set of sterms (that has been compiled from a
rule set of terms). Hence, the correctness theorem only deals with the opposite
It remains to construct a suitable rs such that applying value
to sterm to it
yields the given sterm rule set. We can exploit the side condition (Sect. 5.1) that
all bindings define functions, not constants:
Definition 6 (Global clause set). The mapping global
css :: string ((term×
sterm) list) is obtained by stripping the Sabs constructors from all definitions and
converting the resulting list to a mapping.
For each definition with name f we define a corresponding term v
= Vrecabs
css f []. In other words, each function is now represented by a recursive
closure bundling all functions. Applying value
to sterm to v
returns the original
definition of f.Letrs denote the original sterm rule set and rs
the environment
mapping all f’s to the v
The variable environments σ and σ
can safely be set to the empty mapping,
because top-level terms are evaluated without any free variable bindings.
Corollary 1 (Correctness). rs
, [] t v rs, [] t value to sterm v
Note that this step was not part of the compiler (although rs
is computable)
but it is a refinement of the semantics to support a more modular correctness
Example. Recall the odd and even example from Sect. 4.5. After compilation to
sterm, the rule set looks like this:
rs = {("odd", Sabs [0⇒False , Suc n⇒even n]),
("even", Sabs [0⇒True , Suc n⇒odd n])}
This can be easily transformed into the following global clause set:
css =["odd" → [0⇒False , Suc n⇒even n],
"even" → [0⇒True , Suc n⇒odd n]]
Finally, rs
is computed by creating a recursive closure for each function:
=["odd" → Vrecabs global css "odd" [],
"even" → Vrecabs global
css "even" []]
A Verified Compiler from Isabelle/HOL to CakeML 1017
name / constructors σ name = Some v
σ Sconst name v
Va r
σ name = Some v
σ Svar name v
σ Λ cs Vabs cs σ
σ t Vabs cs σ
σ u v first match cs v = Some (σ
, rhs) σ
++ σ
rhs v
σ t $ u v
σ t Vrecabs css name σ
css name = Some cs σ u v first match cs v = Some (σ
, rhs)
++ mk rec env css σ
++ σ
rhs v
σ t $ u v
name constructors σ t
··· σ t
σ Sconst name $ t
$ ...$ t
Vconstr name [v
Fig. 8. ML-style evaluation semantics
5.7 Evaluation with Recursive Closures
CakeML distinguishes between non-recursive and recursive closures [30]. This
distinction is also present in the value type. In this step, we will conflate vari-
ables with constants which necessitates a special treatment of recursive closures.
Therefore we introduce a new predicate σ t v in Fig. 8 (in contrast to the
previous rs t v). We examine the rules one by one:
Const/Var. Constant definition and variable values are both retrieved from
the same environment σ. We have opted to keep the distinction between
constants and variables in the sterm type to avoid the introduction of another
term type.
Abs. Identical to the previous evaluation semantics. Note that evaluation never
creates recursive closures at run-time (only at compile-time, see Sect. 5.6).
Anonymous functions, e.g. in the term map (λx. x), are evaluated to non-
recursive closures.
Comb. Identical to the previous evaluation semantics.
RecComb. Almost identical to the evaluation semantics. Additionally, for each
function (name, cs) css, a new recursive closure Vrecabs css name σ
created and inserted into the environment. This ensures that after the first
call to a recursive function, the function itself is present in the environment to
be called recursively, without having to introduce coinductive environments.
Constr. Identical to the evaluation semantics.
Conflating Constants and Variables. By merging the rule set rs with the
variable environment σ, it becomes necessary to discuss possible clashes. Previ-
ously, the syntactic distinction between Svar and Sconst meant that x and x
are not ambiguous: all semantics up to the evaluation semantics clearly specify
1018 L. Hupel and T. Nipkow
where to look for the substitute. This is not the case in functional languages
where functions and variables are not distinguished syntactically.
Instead, we rely on the fact that the initial rule set only defines constants. All
variables are introduced by matching before β-reduction (that is, in the Comb
and RecComb rules). The Abs rule does not change the environment. Hence
it suffices to assume that variables in patterns must not overlap with constant
names (see Sect. 5.1).
Correspondence Relation. Both constant definitions and values of variables
are recorded in a single environment σ. This also applies to the environment
contained in a closure. The correspondence relation thus needs to take a different
sets of bindings in closures into account.
Hence, we define a relation
that is implicitly parametrized on the rule
set rs and compares environments. We call it right-conflating, because in a cor-
respondence v
u, any bound environment in u is thought to contain both
variables and constants, whereas in v, any bound environment contains only
Definition 7 (Right-conflating correspondence). We define
tively as follows:
··· v
Vconstr name [v
Vconstr name [u
x frees cs
x x consts cs. rs x
Vabs cs σ
Vabs cs σ
cs range css. x frees cs
cs range css. x consts cs
++ mk rec env css σ
) x
Vrecabs css name σ
Vrecabs css name σ
is not reflexive.
Correctness. The correctness lemma is straightforward to state:
Theorem 4 (Correctness). Let σ be an environment, t beaclosedtermand
v a value such that σ t v. If for all constants x occurring in t,rsx
holds, then there is an u such that rs, [] t u and u
As usual, the rather technical proof proceeds via induction over the semantics
(Fig. 8). It is important to note that the global clause set construction (Sect. 5.6)
satisfies the preconditions of this theorem:
Lemma 4. If name is the name of a constant in rs, then
Vrecabs global
css name []
Vrecabs global css name []
is defined coinductively, the proof of this precondition proceeds by
A Verified Compiler from Isabelle/HOL to CakeML 1019
5.8 CakeML
CakeML is a verified implementation of a subset of Standard ML [24,40]. It
comprises a parser, type checker, formal semantics and backend for machine
code. The semantics has been formalized in Lem [29], which allows export to
Isabelle theories.
Our compiler targets CakeML’s abstract syntax tree. However, we do not
make use of certain CakeML features; notably mutable cells, modules, and lit-
erals. We have derived a smaller, executable version of the original CakeML
semantics, called CupCakeML, together with an equivalence proof. The correct-
ness proof of the last compiler phase establishes a correspondence between Cup-
CakeML and the final semantics of our compiler pipeline.
For the correctness proof of the CakeML compiler, its authors have extracted
the Lem specification into HOL4 theories [1]. In our work, we directly target
CakeML abstract syntax trees (thereby bypassing the parser) and use its big-
step semantics, which we have extracted into Isabelle.
Conversion from sterm to exp. After the series of translations described in the
earlier sections, our terms are syntactically close to CakeML’s terms (Cake.exp).
The only remaining differences are outlined below:
CakeML does not combine abstraction and pattern matching. For that reason,
we have to translate Λ [p
,...]intoΛx. case x of p
| ..., where x
is a fresh variable name. We reuse the fresh monad to obtain a bound variable
name. Note that it is not necessary to thread through already created variable
names, only existing names. The reason is simple: a generated variable is
bound and then immediately used in the body. Shadowing it somewhere in
the body is not problematic.
CakeML has two distinct syntactic categories for identifiers (that can repre-
sent variables or functions) and data constructors. Our term types however
have two distinct syntactic categories for constants (that can represent func-
tions or data constructors) and variables. The necessary prerequisites to deal
with this are already present in the ML-style evaluation semantics (Sect. 5.7)
which conflates constants and variables, but has a dedicated Constr rule for
data constructors.
Types. During embedding (Sect. 3), all type information is erased. Yet, CakeML
performs some limited form of type checking at run-time: constructing and
matching data must always be fully applied. That is, data constructors must
always occur with all arguments supplied on right-hand and left-hand sides.
Fully applied constructors in terms can be easily guaranteed by simple pre-
processing. For patterns however, this must be ensured throughout the com-
pilation pipeline; it is (like other syntactic constraints) another side condition
imposed on the rule set (Sect. 5.1).
Based on a repository snapshot from March 27, 2017 (0c48672).
1020 L. Hupel and T. Nipkow
The shape of datatypes and constructors is managed in CakeML’s environ-
ment. This particular piece of information is allowed to vary in closures, since
ML supports local type definitions. Tracking this would greatly complicate our
proofs. Hence, we fix a global set of constructors and enforce that all values use
exactly that one.
Correspondence R elation. We define two different correspondence relations:
One for values and one for expressions.
Definition 8 (Expression correspondence)
rel e (Svar n)(Cake.Var n)
n/ constructors
rel e (Sconst n)(Cake.Var n)
n constructors rel e t
rel e (Sconst name $ t
$ ...$ t
)(Cake.Con (Some (Cake.Short name)[u
rel e t
rel e t
rel e t
$ t
Cake.App Cake.Opapp [u
n/ ids (Λ [p
,...]) constructors
= mk ml pat p
rel e t
rel e (Λ [p
,...]) (Cake.Fun n (Cake.Mat (Cake.Var n)) [q
rel e tu q
= mk ml pat p
rel e t
rel e (Λ [p
,...]$t)(Cake.Mat u [q
We will explain each of the rules briefly here.
Var. Variables are directly related by identical name.
Const. As described earlier, constructors are treated specially in CakeML. In
order to not confuse functions or variables with data constructors themselves,
we require that the constant name is not a constructor.
Constr. Constructors are directly related by identical name, and recursively
related arguments.
App. CakeML does not just support general function application but also unary
and binary operators. In fact, function application is the binary operator
Opapp. We never generate other operators. Hence the correspondence is
restricted to Opapp.
Fun/Mat. Observe the symmetry between these two cases: In our term lan-
guage, matching and abstraction are combined, which is not the case in
CakeML. This means we relate a case abstraction to a CakeML function con-
taining a match, and a case abstraction applied to a value to just a CakeML
There is no separate relation for patterns, because their translation is simple.
The value correspondence (rel
v) is structurally simpler. In the case of con-
structor values (Vconstr and Cake.Conv), arguments are compared recursively.
Closures and recursive closures are compared extensionally, i.e. only bindings
that occur in the body are checked recursively for correspondence.
A Verified Compiler from Isabelle/HOL to CakeML 1021
Correctness. We use the same trick as in Sect. 5.6 to obtain a suitable envi-
ronment for CakeML evaluation based on the rule set rs.
Theorem 5 ( Correctness). If the compiled expression sterm
to cake t termi-
nates with a value u in the CakeML semantics, there is a value v such that
v vuand rs t v.
6 Composition
The complete compiler pipeline consists of multiple phases. Correctness is justi-
fied for each phase between intermediate semantics and correspondence relations,
most of which are rather technical. Whereas the compiler may be complex and
impenetrable, the trustworthiness of the constructions hinges on the obviousness
of those correspondence relations.
Fortunately, under the assumption that terms to be evaluated and the result-
ing values do not contain abstractions or closures, respectively all of the
correspondence relations collapse to simple structural equality: two terms are
related if and only if one can be converted to the other by consistent renaming
of term constructors.
The actual compiler can be characterized with two functions. Firstly, the
translation of term to Cake.exp is a simple composition of each term translation
definition term to cake :: term Cake.exp where
to cake = s term to cake pterm to sterm nterm to pterm term to nterm
Secondly, the function that translates function definitions by composing the
phases as outlined in Fig. 2, including iterated application of pattern elimination:
definition compile :: (term × term) fset Cake.dec where
compile = Cake.Dletrec compile
srules to cake compile prules to srules
irules to srules compile irules iter compile crules to irules
of compile rules to nrules
Each function compile * corresponds to one compiler phase; the remaining func-
tions are trivial. This produces a CakeML top-level declaration. We prove that
evaluating this declaration in the top-level semantics (evaluate
prog) results in an
environment cake
sem env.Butcake sem env can also be computed via another
instance of the global clause set trick (Sect. 5.6).
Equipped with these functions, we can state the final correctness theorem:
theorem compiled correct:
(If CakeML evaluation of a term succeeds ... )
assumes evaluate False cake
sem env s (term to cake t)(s’, Rval ml v)
(... producing a constructor term without closures ... )
assumes cake
abstraction free ml v
(... and some syntactic properties of the involved terms hold ... )
assumes closed t and ¬ shadows
consts (heads rs constructors) t and
welldefined (heads rs constructors) t and wellformed t
(... then this evaluation can be reproduced in the termrewriting semantics )
shows rs t
cake to term ml v
1022 L. Hupel and T. Nipkow
class add =
fixes plus :: ’a ’a ’a
definition f :: (’a::add) ’a where
f x = plus xx
(a) Source program
datatype ’a dict
add = Dict add (’a ’a ’a)
fun cert
add :: (’a::add) dict add bool where
add (Dict add pls)=(pls =plus)
fun f’ :: ’a dict
add ’a ’a where
f’ (Dict
add pls) x = pls x x
lemma f’
eq: cert add dict f’ dict =f
(b) Result of translation
Fig. 9. Dictionary construction in Isabelle
This theorem directly relates the evaluation of a term t in the full CakeML
(including mutability and exceptions) to the evaluation in the initial higher-order
term rewriting semantics. The evaluation of t happens using the environment
produced from the initial rule set. Hence, the theorem can be interpreted as the
correctness of the pseudo-ML expression let rec rs in t.
Observe that in the assumption, the conversion goes from our terms to
CakeML expressions, whereas in the conclusion, the conversion goes the opposite
7 Dictionary Construction
Isabelle’s type system supports type classes (or simply classes)[18,44] whereas
CakeML does not. In order to not complicate the correctness proofs, type classes
are not supported by our embedded term language either. Instead, we eliminate
classes and instances by a dictionary construction [19] before embedding into the
term language. Haftmann and Nipkow give a pen-and-paper correctness proof
of this construction [17, Sect. 4.1]. We augmented the dictionary construction
with the generation of a certificate theorem that shows the equivalence of the
two versions of a function, with type classes and with dictionaries. This section
briefly explains our dictionary construction.
Figure 9 shows a simple example of a dictionary construction. Type vari-
ables may carry class constraints (e.g. α :: add). The basic idea is that classes
become dictionaries containing the functions of that class; class instances become
dictionary definitions. Dictionaries are realized as datatypes. Class constraints
become additional dictionary parameters for that class. In the example, class
add becomes dict
add; function f is translated into f
which takes an additional
parameter of type dict
add. In reality our tool does not produce the Isabelle
source code shown in Fig. 9b but performs the constructions internally. The cor-
rectness lemma f
eq is proved automatically. Its precondition expresses that the
dictionary must contain exactly the function(s) of class add. For any monomor-
phic instance, the precondition can be proved outright based on the certificate
theorems proved for each class instance as explained next.
A Verified Compiler from Isabelle/HOL to CakeML 1023
Not shown in the example is the translation of class instances. The basic
form of a class instance in Isabelle is τ::(c
) c where τ is an n-ary type
constructor. It corresponds to Haskell’s (c
) c (τα
and is translated into a function inst
c τ ::α
dict c
··· α
dict c
) τ dict c and the following certificate theorem is proved:
···cert c
cert c (inst c τ dict
... dict
For a more detailed explanation of how the dictionary construction works, we
refer to the corresponding entry in the Archive of Formal Proofs [21].
8 Evaluation
We have tried out our compiler on examples from existing Isabelle formalizations.
This includes an implementation of Huffman encoding, lists and sorting, string
functions [39], and various data structures from Okasaki’s book [34], including
binary search trees, pairing heaps, and leftist heaps. These definitions can be
processed with slight modifications: functions need to be totalized (see the end
of Sect. 3). However, parts of the tactics required for deep embedding proofs
(Sect. 3) are too slow on some functions and hence still need to be optimized.
9 Conclusion
For this paper we have concentrated on the compiler from Isabelle/HOL to
CakeML abstract syntax trees. Partial correctness is proved w.r.t. the big-step
semantics of CakeML. In the next step we will link our work with the compiler
from CakeML to machine code. Tan et al. [40, Sect. 10] prove a correctness the-
orem that relates their semantics with the execution of the compiled machine
code. In that paper, they use a newer iteration of the CakeML semantics (func-
tional big-step [35]) than we do here. Both semantics are still present in the
CakeML source repository, together with an equivalence proof. Another impor-
tant step consists of targeting CakeML’s native types, e.g. integer numbers and
Evaluation of our compiled programs is already possible via Isabelle’s pred-
icate compiler [5], which allows us to turn CakeML’s big-step semantics into
an executable function. We have used this execution mechanism to establish for
sample programs that they terminate successfully. We also plan to prove that
our compiled programs terminate, i.e. total correctness.
The total size of this formalization, excluding theories extracted from Lem,
is currently approximately 20000 lines of proof text (90 %) and ML code (10 %).
The ML code itself produces relatively simple theorems, which means that there
are less opportunities for it to go wrong. This constitutes an improvement over
certifying approaches that prove complicated properties in ML.
1024 L. Hupel and T. Nipkow
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