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We have implemented the algorithmic system $\vdashA$. Our
implementation is written in OCaml and uses CDuce as a library to
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provide the semantic subtyping machinery. Besides a type-checking
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algorithm defined on the base language, our implementation supports
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record types (Section \ref{ssec:struct}) and the refinement of
function types (Section \ref{sec:refining} with the rule of
Appendix~\ref{app:optimize}). The implementation is rather crude and
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consists of 2000 lines of OCaml code, including parsing, type-checking
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of programs, and pretty printing of types. We demonstrate the output
of our type-checking implementation in Table~\ref{tab:implem} by
listing some examples none of which can be typed by current
systems (even though some system such as Flow and TypeScript
can type some of them by adding explicit type annotation, the code 6,
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7, 9, and 10 in Table~\ref{tab:implem}  and, even more,  the \code{and\_} and \code{xor\_} functions at the end of this
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section are out of reach of current systems, even when using the right
explicit annotations). These examples and others can be tested in the
online toplevel available at
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\url{https://occtyping.github.io/}%
\ifsubmission
~(the corresponding repository is
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anonymized).
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\else.
\fi
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\input{code_table}
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In this table, the second column gives a code fragment and the third
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column the type deduced by our implementation. Code~1 is a
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straightforward function similar to our introductory example \code{foo} in (\ref{foo},\ref{foo2}). Here the
programmer annotates the parameter of the function with a coarse type
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$\Int\vee\Bool$. Our implementation first type-checks the body of the
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function under this assumption, but doing so collects that the type of
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$\texttt{x}$ is specialized to \Int{} in the ``then'' case and to \Bool{}
in the ``else'' case. The function is thus type-checked twice more
under each hypothesis for \texttt{x}, yielding the precise type
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$(\Int\to\Int)\land(\Bool\to\Bool)$. Note that w.r.t.\  rule \Rule{AbsInf+} of Section~\ref{sec:refining}, our implementation improved the output of the computed
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type. Indeed using rule~[{\sc AbsInf}+] we would obtain the
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type
$(\Int\to\Int)\land(\Bool\to\Bool)\land(\Bool\vee\Int\to\Bool\vee\Int)$
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with a redundant arrow. Here we can see that since we deduced
the first two arrows $(\Int\to\Int)\land(\Bool\to\Bool)$, and since
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the union of their domain exactly covers the domain of the third arrow,
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the latter is not needed. Code~2 shows what happens when the argument
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of the function is left unannotated (i.e., it is annotated by the top
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type \Any, written ``\texttt{Any}'' in our implementation). Here
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type-checking and refinement also work as expected, but the function
only type checks if all cases for \texttt{x} are covered (which means
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that the function must handle the case of inputs that are neither in \Int{}
nor in \Bool).
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The following examples paint a more interesting picture. First
(Code~3) it is
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easy in our formalism to program type predicates such as those
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hard-coded in the $\lambda_{\textit{TR}}$ language of \citet{THF10}. Such type
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predicates, which return \texttt{true} if and only if their input has
a particular type, are just plain functions with an intersection
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type inferred by the system of Section~\ref{sec:refining}. We next define Boolean connectives as overloaded
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functions. The \texttt{not\_} connective (Code~4) just tests whether its
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argument is the Boolean \texttt{true} by testing that it belongs to
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the singleton type \True{} (the type whose only value is
\texttt{true}) returning \texttt{false} for it and \texttt{true} for
any other value (recall that $\neg\True$ is equivalent to
$\texttt{Any\textbackslash}\True$). It works on values of any type,
but we could restrict it to Boolean values by simply annotating the
parameter by \Bool{} (which in CDuce is syntactic sugar for
\True$\vee$\False) yielding the type
$(\True{\to}\False)\wedge(\False{\to}\True)$.
The \texttt{or\_} connective (Code~5) is straightforward as far as the
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code goes, but we see that the overloaded type precisely captures all
possible cases. Again we use a generalized version of the
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\texttt{or\_} connective that accepts and treats any value that is not
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\texttt{true} as \texttt{false} and again, we could easily restrict the
domain to \Bool{} if desired.
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To showcase the power of our type system, and in particular of
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the ``$\worra{}{}$''
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type operator, we define \texttt{and\_} (Code~6) using De Morgan's
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Laws instead of
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using a direct definition. Here the application of the outermost \texttt{not\_} operator is checked against type \True. This
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allows the system to deduce that the whole \texttt{or\_} application
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has type \False, which in turn leads to \texttt{not\_\;x} and
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\texttt{not\_\;y} to have type $\lnot \True$ and therefore both \texttt{x}
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and \texttt{y} to have type \True. The whole function is typed with
the most precise type (we present the type as printed by our
implementation, but the first arrow of the resulting type is
equivalent to
$(\True\to\lnot\True\to\False)\land(\True\to\True\to\True)$).

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All these type predicates and Boolean connectives can be used together
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to write complex type tests, as in Code~7. Here we define a function
\texttt{f} that takes two arguments \texttt{x} and \texttt{y}. If
\texttt{x} is an integer and \texttt{y} a Boolean, then it returns the
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integer \texttt{1}; if \texttt{x} is a character or
\texttt{y} is an integer, then it returns \texttt{2}; otherwise the
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function returns \texttt{3}. Our system correctly deduces a (complex)
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intersection type that covers all cases (plus several redundant arrow
types). That this type is as precise as possible can be shown by the fact that
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when applying
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\texttt{f} to arguments of the expected type, the \emph{type} deduced for the
whole expression is the singleton type \texttt{1}, or \texttt{2},
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or \texttt{3}, depending on the type of the arguments.

Code~8 allows us to demonstrate the use and typing of record paths. We
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model, using open records, the type of DOM objects that represent XML
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or HTML documents. Such objects possess a common field
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\texttt{nodeType} containing an integer constant denoting the kind of
the node (e.g., \p{9} for the root element, \p{1} for an element node, \p{3} for a text node, \ldots). Depending on the kind, the object will have
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different fields and methods. It is common practice to perform a test
on the value of the \texttt{nodeType} field. In dynamic languages such
as JavaScript, the relevant field or method can directly be accessed
after having checked for the appropriate \texttt{nodeType}. In
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mainstream statically typed languages, such as Java, a downward cast
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from the generic \texttt{Node} type to the expected precise type of
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the object is needed. We can see that using the extension presented in
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Section~\ref{ssec:struct} we can deduce the correct type for
\texttt{x} in all cases. Of particular interest is the last case,
since we use a type case to check the emptiness of the list of child
nodes. This splits, at the type level, the case for the \Keyw{Element}
type depending on whether the content of the \texttt{childNodes} field
is the empty list or not.

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Code~9 shows the usefulness of the rule \Rule{OverApp}.
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Consider the definition of the \texttt{xor\_} operator.
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Here the rule~[{\sc AbsInf}+] is not sufficient to precisely type the
function, and using only this rule would yield a type
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 $\Any\to\Any\to\Bool$.
\iflongversion
Let us follow the behavior of the
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 ``$\worra{}{}$'' operator. Here the whole \texttt{and\_} is requested
 to have type \True, which implies that \texttt{or\_ x y} must have
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 type \True. This can always happen, whether \texttt{x} is \True{} or
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 not (but then depends on the type of \texttt{y}). The ``$\worra{}{}$''
 operator correctly computes that the type for \texttt{x} in the
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 ``\texttt{then}'' branch is $\True\vee\lnot\True\lor\True\simeq\Any$,
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and a similar reasoning holds for \texttt{y}. 
\fi%%%%%%%%%%%%%%
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However, since \texttt{or\_} has type
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%\\[.7mm]\centerline{%
$(\True\to\Any\to\True)\land(\Any\to\True\to\True)\land
   (\lnot\True\to\lnot\True\to\False)$
%}\\[.7mm]
then the rule \Rule{OverApp} applies and \True, \Any, and $\lnot\True$ become candidate types for
\texttt{x}, which allows us to deduce the precise type given in the table. Finally, thanks to rule \Rule{OverApp} it is not  necessary to use a type case to force refinement. As a consequence we can define the functions \texttt{and\_} and \texttt{xor\_} more naturally as:
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\begin{alltt}\color{darkblue}\morecompact
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  let and_ = fun (x : Any) -> fun (y : Any) -> not_ (or_ (not_ x) (not_ y))
  let xor_ = fun (x : Any) -> fun (y : Any) -> and_ (or_ x y) (not_ (and_ x y))
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\end{alltt}
for which the very same types as in Table~\ref{tab:implem} are deduced.
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Last but not least Code~10 (corresponding to our introductory
example~\eqref{nest1}) illustrates the need for iterative refinement of
type environments, as defined in Section~\ref{sec:typenv}. As
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explained, a single pass analysis would deduce 
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for {\tt x}
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a type \Int{} from the {\tt f\;x} application and \Any{} from the {\tt g\;x}
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application. Here by iterating a second time, the algorithm deduces
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that {\tt x} has type $\Empty$ (i.e., $\textsf{Empty}$), that is that the first branch can never
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be selected (and our implementation warns the user accordingly). In hindsight, the only way for a well-typed overloaded function to have
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type $(\Int{\to}\Int)\land(\Any{\to}\Bool)$ is to diverge when the
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argument is of type \Int: since this intersection type states that
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whenever the input is \Int, {\em both\/} branches can be selected,
yielding a result that is at the same time an integer and a Boolean.
This is precisely reflected by the case $\Int\to\Empty$ in the result.
Indeed our {\tt example10} function can be applied to an integer, but
at runtime the application of {\tt f ~x} will diverge.